[PATCHSET v2 00/14] xfs: design documentation for online fsck

[PATCHSET v2 00/14] xfs: design documentation for online fsck

[Darrick J. Wong]

Hi all,

To prepare the XFS community and potential patch reviewers for the
upstream submission of the online fsck feature, I decided to write a
document capturing the broader picture behind the online repair
development effort.  The document begins by defining the problems that
online fsck aims to solve and outlining specific use cases for the
functionality.

Using that as a base, the rest of the design document presents the high
level algorithms that fulfill the goals set out at the start and the
interactions between the large pieces of the system.  Case studies round
out the design documentation by adding the details of exactly how
specific parts of the online fsck code integrate the algorithms with the
filesystem.

The goal of this effort is to help the XFS community understand how the
gigantic online repair patchset works.  The questions I submit to the
community reviewers are:

1. As you read the design doc (and later the code), do you feel that you
   understand what's going on well enough to try to fix a bug if you
   found one?

2. What sorts of interactions between systems (or between scrub and the
   rest of the kernel) am I missing?

3. Do you feel confident enough in the implementation as it is now that
   the benefits of merging the feature (as EXPERIMENTAL) outweigh any
   potential disruptions to XFS at large?

4. Are there problematic interactions between subsystems that ought to
   be cleared up before merging?

I intend to commit this document to the kernel's documentation directory
around the time we start merging the patchset, albeit without the links
to git.kernel.org.  A much more readable version of this is posted at:
https://djwong.org/docs/xfs-online-fsck-design/

v2: add missing sections about: all the in-kernel data structures and
    new apis that the scrub and repair functions use; how xattrs and
    directories are checked; how space btree records are checked; and
    add more details to the parts where all these bits tie together.
    Proofread for verb tense inconsistencies and eliminate vague 'we'
    usage.  Move all the discussion of what we can do with pageable
    kernel memory into a single source file and section.  Document where
    log incompat feature locks fit into the locking model.

If you're going to start using this mess, you probably ought to just
pull from my git trees, which are linked below.

This is an extraordinary way to destroy everything.  Enjoy!
Comments and questions are, as always, welcome.

--D

kernel git tree:
https://git.kernel.org/cgit/linux/kernel/git/djwong/xfs-linux.git/log/?h=online-fsck-design
---
 Documentation/filesystems/index.rst                |    1 
 .../filesystems/xfs-online-fsck-design.rst         | 4979 ++++++++++++++++++++
 .../filesystems/xfs-self-describing-metadata.rst   |    1 
 3 files changed, 4981 insertions(+)
 create mode 100644 Documentation/filesystems/xfs-online-fsck-design.rst

[PATCH 01/14] xfs: document the motivation for online fsck design

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Start the first chapter of the online fsck design documentation.
This covers the motivations for creating this in the first place.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 Documentation/filesystems/index.rst                |    1 
 .../filesystems/xfs-online-fsck-design.rst         |  199 ++++++++++++++++++++
 2 files changed, 200 insertions(+)
 create mode 100644 Documentation/filesystems/xfs-online-fsck-design.rst


diff --git a/Documentation/filesystems/index.rst b/Documentation/filesystems/index.rst
index bee63d42e5ec..fbb2b5ada95b 100644
--- a/Documentation/filesystems/index.rst
+++ b/Documentation/filesystems/index.rst
@@ -123,4 +123,5 @@ Documentation for filesystem implementations.
    vfat
    xfs-delayed-logging-design
    xfs-self-describing-metadata
+   xfs-online-fsck-design
    zonefs
diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
new file mode 100644
index 000000000000..25717ebb5f80
--- /dev/null
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -0,0 +1,199 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. _xfs_online_fsck_design:
+
+..
+        Mapping of heading styles within this document:
+        Heading 1 uses "====" above and below
+        Heading 2 uses "===="
+        Heading 3 uses "----"
+        Heading 4 uses "````"
+        Heading 5 uses "^^^^"
+        Heading 6 uses "~~~~"
+        Heading 7 uses "...."
+
+        Sections are manually numbered because apparently that's what everyone
+        does in the kernel.
+
+======================
+XFS Online Fsck Design
+======================
+
+This document captures the design of the online filesystem check feature for
+XFS.
+The purpose of this document is threefold:
+
+- To help kernel distributors understand exactly what the XFS online fsck
+  feature is, and issues about which they should be aware.
+
+- To help people reading the code to familiarize themselves with the relevant
+  concepts and design points before they start digging into the code.
+
+- To help developers maintaining the system by capturing the reasons
+  supporting higher level decisionmaking.
+
+As the online fsck code is merged, the links in this document to topic branches
+will be replaced with links to code.
+
+This document is licensed under the terms of the GNU Public License, v2.
+The primary author is Darrick J. Wong.
+
+This design document is split into seven parts.
+Part 1 defines what fsck tools are and the motivations for writing a new one.
+Parts 2 and 3 present a high level overview of how online fsck process works
+and how it is tested to ensure correct functionality.
+Part 4 discusses the user interface and the intended usage modes of the new
+program.
+Parts 5 and 6 show off the high level components and how they fit together, and
+then present case studies of how each repair function actually works.
+Part 7 sums up what has been discussed so far and speculates about what else
+might be built atop online fsck.
+
+.. contents:: Table of Contents
+   :local:
+
+1. What is a Filesystem Check?
+==============================
+
+A Unix filesystem has three main jobs: to provide a hierarchy of names through
+which application programs can associate arbitrary blobs of data for any
+length of time, to virtualize physical storage media across those names, and
+to retrieve the named data blobs at any time.
+The filesystem check (fsck) tool examines all the metadata in a filesystem
+to look for errors.
+Simple tools only check for obvious corruptions, but the more sophisticated
+ones cross-reference metadata records to look for inconsistencies.
+People do not like losing data, so most fsck tools also contains some ability
+to deal with any problems found.
+As a word of caution -- the primary goal of most Linux fsck tools is to restore
+the filesystem metadata to a consistent state, not to maximize the data
+recovered.
+That precedent will not be challenged here.
+
+Filesystems of the 20th century generally lacked any redundancy in the ondisk
+format, which means that fsck can only respond to errors by erasing files until
+errors are no longer detected.
+System administrators avoid data loss by increasing the number of separate
+storage systems through the creation of backups; and they avoid downtime by
+increasing the redundancy of each storage system through the creation of RAID.
+More recent filesystem designs contain enough redundancy in their metadata that
+it is now possible to regenerate data structures when non-catastrophic errors
+occur; this capability aids both strategies.
+Over the past few years, XFS has added a storage space reverse mapping index to
+make it easy to find which files or metadata objects think they own a
+particular range of storage.
+Efforts are under way to develop a similar reverse mapping index for the naming
+hierarchy, which will involve storing directory parent pointers in each file.
+With these two pieces in place, XFS uses secondary information to perform more
+sophisticated repairs.
+
+TLDR; Show Me the Code!
+-----------------------
+
+Code is posted to the kernel.org git trees as follows:
+`kernel changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-symlink>`_,
+`userspace changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-media-scan-service>`_, and
+`QA test changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=repair-dirs>`_.
+Each kernel patchset adding an online repair function will use the same branch
+name across the kernel, xfsprogs, and fstests git repos.
+
+Existing Tools
+--------------
+
+The online fsck tool described here will be the third tool in the history of
+XFS (on Linux) to check and repair filesystems.
+Two programs precede it:
+
+The first program, ``xfs_check``, was created as part of the XFS debugger
+(``xfs_db``) and can only be used with unmounted filesystems.
+It walks all metadata in the filesystem looking for inconsistencies in the
+metadata, though it lacks any ability to repair what it finds.
+Due to its high memory requirements and inability to repair things, this
+program is now deprecated and will not be discussed further.
+
+The second program, ``xfs_repair``, was created to be faster and more robust
+than the first program.
+Like its predecessor, it can only be used with unmounted filesystems.
+It uses extent-based in-memory data structures to reduce memory consumption,
+and tries to schedule readahead IO appropriately to reduce I/O waiting time
+while it scans the metadata of the entire filesystem.
+The most important feature of this tool is its ability to respond to
+inconsistencies in file metadata and directory tree by erasing things as needed
+to eliminate problems.
+Space usage metadata are rebuilt from the observed file metadata.
+
+Problem Statement
+-----------------
+
+The current XFS tools leave several problems unsolved:
+
+1. **User programs** suddenly **lose access** to information in the computer
+   when unexpected shutdowns occur as a result of silent corruptions in the
+   filesystem metadata.
+   These occur **unpredictably** and often without warning.
+
+2. **Users** experience a **total loss of service** during the recovery period
+   after an **unexpected shutdown** occurs.
+
+3. **Users** experience a **total loss of service** if the filesystem is taken
+   offline to **look for problems** proactively.
+
+4. **Data owners** cannot **check the integrity** of their stored data without
+   reading all of it.
+   This may expose them to substantial billing costs when a linear media scan
+   might suffice.
+
+5. **System administrators** cannot **schedule** a maintenance window to deal
+   with corruptions if they **lack the means** to assess filesystem health
+   while the filesystem is online.
+
+6. **Fleet monitoring tools** cannot **automate periodic checks** of filesystem
+   health when doing so requires **manual intervention** and downtime.
+
+7. **Users** can be tricked into **doing things they do not desire** when
+   malicious actors **exploit quirks of Unicode** to place misleading names
+   in directories.
+
+Given this definition of the problems to be solved and the actors who would
+benefit, the proposed solution is a third fsck tool that acts on a running
+filesystem.
+
+This new third program has three components: an in-kernel facility to check
+metadata, an in-kernel facility to repair metadata, and a userspace driver
+program to drive fsck activity on a live filesystem.
+``xfs_scrub`` is the name of the driver program.
+The rest of this document presents the goals and use cases of the new fsck
+tool, describes its major design points in connection to those goals, and
+discusses the similarities and differences with existing tools.
+
++--------------------------------------------------------------------------+
+| **Note**:                                                                |
++--------------------------------------------------------------------------+
+| Throughout this document, the existing offline fsck tool can also be     |
+| referred to by its current name "``xfs_repair``".                        |
+| The userspace driver program for the new online fsck tool can be         |
+| referred to as "``xfs_scrub``".                                          |
+| The kernel portion of online fsck that validates metadata is called      |
+| "online scrub", and portion of the kernel that fixes metadata is called  |
+| "online repair".                                                         |
++--------------------------------------------------------------------------+
+
+Secondary metadata indices enable the reconstruction of parts of a damaged
+primary metadata object from secondary information.
+XFS filesystems shard themselves into multiple primary objects to enable better
+performance on highly threaded systems and to contain the blast radius when
+problems happen.
+The naming hierarchy is broken up into objects known as directories and files;
+and the physical space is split into pieces known as allocation groups.
+The division of the filesystem into principal objects (allocation groups and
+inodes) means that there are ample opportunities to perform targeted checks and
+repairs on a subset of the filesystem.
+While this is going on, other parts continue processing IO requests.
+Even if a piece of filesystem metadata can only be regenerated by scanning the
+entire system, the scan can still be done in the background while other file
+operations continue.
+
+In summary, online fsck takes advantage of resource sharding and redundant
+metadata to enable targeted checking and repair operations while the system
+is running.
+This capability will be coupled to automatic system management so that
+autonomous self-healing of XFS maximizes service availability.

[PATCH 02/14] xfs: document the general theory underlying online fsck design

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Start the second chapter of the online fsck design documentation.
This covers the general theory underlying how online fsck works.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  366 ++++++++++++++++++++
 1 file changed, 366 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 25717ebb5f80..a03a7b9f0250 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -197,3 +197,369 @@ metadata to enable targeted checking and repair operations while the system
 is running.
 This capability will be coupled to automatic system management so that
 autonomous self-healing of XFS maximizes service availability.
+
+2. Theory of Operation
+======================
+
+Because it is necessary for online fsck to lock and scan live metadata objects,
+online fsck consists of three separate code components.
+The first is the userspace driver program ``xfs_scrub``, which is responsible
+for identifying individual metadata items, scheduling work items for them,
+reacting to the outcomes appropriately, and reporting results to the system
+administrator.
+The second and third are in the kernel, which implements functions to check
+and repair each type of online fsck work item.
+
++------------------------------------------------------------------+
+| **Note**:                                                        |
++------------------------------------------------------------------+
+| For brevity, this document shortens the phrase "online fsck work |
+| item" to "scrub item".                                           |
++------------------------------------------------------------------+
+
+Scrub item types are delineated in a manner consistent with the Unix design
+philosophy, which is to say that each item should handle one aspect of a
+metadata structure, and handle it well.
+
+Scope
+-----
+
+In principle, online fsck should be able to check and to repair everything that
+the offline fsck program can handle.
+However, the adjective *online* brings with it the limitation that online fsck
+cannot deal with anything that prevents the filesystem from going on line, i.e.
+mounting.
+This limitation means that maintenance of the offline fsck tool will continue.
+A second limitation of online fsck is that it must follow the same resource
+sharing and lock acquisition rules as the regular filesystem.
+This means that scrub cannot take *any* shortcuts to save time, because doing
+so could lead to concurrency problems.
+In other words, online fsck will never be able to fix 100% of the
+inconsistencies that offline fsck can repair, and a complete run of online fsck
+may take longer.
+However, both of these limitations are acceptable tradeoffs to satisfy the
+different motivations of online fsck, which are to **minimize system downtime**
+and to **increase predictability of operation**.
+
+.. _scrubphases:
+
+Phases of Work
+--------------
+
+The userspace driver program ``xfs_scrub`` splits the work of checking and
+repairing an entire filesystem into seven phases.
+Each phase concentrates on checking specific types of scrub items and depends
+on the success of all previous phases.
+The seven phases are as follows:
+
+1. Collect geometry information about the mounted filesystem and computer,
+   discover the online fsck capabilities of the kernel, and open the
+   underlying storage devices.
+
+2. Check allocation group metadata, all realtime volume metadata, and all quota
+   files.
+   Each metadata structure is scheduled as a separate scrub item.
+   If corruption is found in the inode header or inode btree and ``xfs_scrub``
+   is permitted to perform repairs, then those scrub items are repaired to
+   prepare for phase 3.
+   Repairs are implemented by resubmitting the scrub item to the kernel with
+   the repair flag enabled; this is discussed in the next section.
+   Optimizations and all other repairs are deferred to phase 4.
+
+3. Check all metadata of every file in the filesystem.
+   Each metadata structure is also scheduled as a separate scrub item.
+   If repairs are needed, ``xfs_scrub`` is permitted to perform repairs,
+   and there were no problems detected during phase 2, then those scrub items
+   are repaired.
+   Optimizations and unsuccessful repairs are deferred to phase 4.
+
+4. All remaining repairs and scheduled optimizations are performed during this
+   phase, if the caller permits them.
+   Before starting repairs, the summary counters are checked and any necessary
+   repairs are performed so that subsequent repairs will not fail the resource
+   reservation step due to wildly incorrect summary counters.
+   Unsuccesful repairs are requeued as long as forward progress on repairs is
+   made somewhere in the filesystem.
+   Free space in the filesystem is trimmed at the end of phase 4 if the
+   filesystem is clean.
+
+5. By the start of this phase, all primary and secondary filesystem metadata
+   must be correct.
+   Summary counters such as the free space counts and quota resource counts
+   are checked and corrected.
+   Directory entry names and extended attribute names are checked for
+   suspicious entries such as control characters or confusing Unicode sequences
+   appearing in names.
+
+6. If the caller asks for a media scan, read all allocated and written data
+   file extents in the filesystem.
+   The ability to use hardware-assisted data file integrity checking is new
+   to online fsck; neither of the previous tools have this capability.
+   If media errors occur, they will be mapped to the owning files and reported.
+
+7. Re-check the summary counters and presents the caller with a summary of
+   space usage and file counts.
+
+Steps for Each Scrub Item
+-------------------------
+
+The kernel scrub code uses a three-step strategy for checking and repairing
+the one aspect of a metadata object represented by a scrub item:
+
+1. The scrub item of interest is checked for corruptions; opportunities for
+   optimization; and for values that are directly controlled by the system
+   administrator but look suspicious.
+   If the item is not corrupt or does not need optimization, resource are
+   released and the positive scan results are returned to userspace.
+   If the item is corrupt or could be optimized but the caller does not permit
+   this, resources are released and the negative scan results are returned to
+   userspace.
+   Otherwise, the kernel moves on to the second step.
+
+2. The repair function is called to rebuild the data structure.
+   Repair functions generally choose rebuild a structure from other metadata
+   rather than try to salvage the existing structure.
+   If the repair fails, the scan results from the first step are returned to
+   userspace.
+   Otherwise, the kernel moves on to the third step.
+
+3. In the third step, the kernel runs the same checks over the new metadata
+   item to assess the efficacy of the repairs.
+   The results of the reassessment are returned to userspace.
+
+Classification of Metadata
+--------------------------
+
+Each type of metadata object (and therefore each type of scrub item) is
+classified as follows:
+
+Primary Metadata
+````````````````
+
+Metadata structures in this category should be most familiar to filesystem
+users either because they are directly created by the user or they index
+objects created by the user
+Most filesystem objects fall into this class.
+Resource and lock acquisition for scrub code follows the same order as regular
+filesystem accesses.
+
+Primary metadata objects are the simplest for scrub to process.
+The principal filesystem object (either an allocation group or an inode) that
+owns the item being scrubbed is locked to guard against concurrent updates.
+The check function examines every record associated with the type for obvious
+errors and cross-references healthy records against other metadata to look for
+inconsistencies.
+Repairs for this class of scrub item are simple, since the repair function
+starts by holding all the resources acquired in the previous step.
+The repair function scans available metadata as needed to record all the
+observations needed to complete the structure.
+Next, it stages the observations in a new ondisk structure and commits it
+atomically to complete the repair.
+Finally, the storage from the old data structure are carefully reaped.
+
+Because ``xfs_scrub`` locks a primary object for the duration of the repair,
+this is effectively an offline repair operation performed on a subset of the
+filesystem.
+This minimizes the complexity of the repair code because it is not necessary to
+handle concurrent updates from other threads, nor is it necessary to access
+any other part of the filesystem.
+As a result, indexed structures can be rebuilt very quickly, and programs
+trying to access the damaged structure will be blocked until repairs complete.
+The only infrastructure needed by the repair code are the staging area for
+observations and a means to write new structures to disk.
+Despite these limitations, the advantage that online repair holds is clear:
+targeted work on individual shards of the filesystem avoids total loss of
+service.
+
+This mechanism is described in section 2.1 ("Off-Line Algorithm") of
+V. Srinivasan and M. J. Carey, `"Performance of On-Line Index Construction
+Algorithms" <https://dl.acm.org/doi/10.5555/645336.649870>`_,
+*Extending Database Technology*, pp. 293-309, 1992.
+
+Most primary metadata repair functions stage their intermediate results in an
+in-memory array prior to formatting the new ondisk structure, which is very
+similar to the list-based algorithm discussed in section 2.3 ("List-Based
+Algorithms") of Srinivasan.
+However, any data structure builder that maintains a resource lock for the
+duration of the repair is *always* an offline algorithm.
+
+Secondary Metadata
+``````````````````
+
+Metadata structures in this category reflect records found in primary metadata,
+but are only needed for online fsck or for reorganization of the filesystem.
+Resource and lock acquisition for scrub code do not follow the same order as
+regular filesystem accesses, and may involve full filesystem scans.
+
+Secondary metadata objects are difficult for scrub to process, because scrub
+attaches to the secondary object but needs to check primary metadata, which
+runs counter to the usual order of resource acquisition.
+Check functions can be limited in scope to reduce runtime.
+Repairs, however, require a full scan of primary metadata, which can take a
+long time to complete.
+Under these conditions, ``xfs_scrub`` cannot lock resources for the entire
+duration of the repair.
+
+Instead, repair functions set up an in-memory staging structure to store
+observations.
+Depending on the requirements of the specific repair function, the staging
+index can have the same format as the ondisk structure, or it can have a design
+specific to that repair function.
+The next step is to release all locks and start the filesystem scan.
+When the repair scanner needs to record an observation, the staging data are
+locked long enough to apply the update.
+Simultaneously, the repair function hooks relevant parts of the filesystem to
+apply updates to the staging data if the the update pertains to an object that
+has already been scanned by the index builder.
+Once the scan is done, the owning object is re-locked, the live data is used to
+write a new ondisk structure, and the repairs are committed atomically.
+The hooks are disabled and the staging staging area is freed.
+Finally, the storage from the old data structure are carefully reaped.
+
+Introducing concurrency helps online repair avoid various locking problems, but
+comes at a high cost to code complexity.
+Live filesystem code has to be hooked so that the repair function can observe
+updates in progress.
+The staging area has to become a fully functional parallel structure so that
+updates can be merged from the hooks.
+Finally, the hook, the filesystem scan, and the inode locking model must be
+sufficiently well integrated that a hook event can decide if a given update
+should be applied to the staging structure.
+
+In theory, the scrub implementation could apply these same techniques for
+primary metadata, but doing so would make it massively more complex and less
+performant.
+Programs attempting to access the damaged structures are not blocked from
+operation, which may cause application failure or an unplanned filesystem
+shutdown.
+
+Inspiration for the secondary metadata repair strategy was drawn from section
+2.4 of Srinivasan above, and sections 2 ("NSF: Inded Build Without Side-File")
+and 3.1.1 ("Duplicate Key Insert Problem") in C. Mohan, `"Algorithms for
+Creating Indexes for Very Large Tables Without Quiescing Updates"
+<https://dl.acm.org/doi/10.1145/130283.130337>`_, 1992.
+
+The sidecar index mentioned above bears some resemblance to the side file
+method mentioned in Srinivasan and Mohan.
+Their method consists of an index builder that extracts relevant record data to
+build the new structure as quickly as possible; and an auxiliary structure that
+captures all updates that would be committed to the index by other threads were
+the new index already online.
+After the index building scan finishes, the updates recorded in the side file
+are applied to the new index.
+To avoid conflicts between the index builder and other writer threads, the
+builder maintains a publicly visible cursor that tracks the progress of the
+scan through the record space.
+To avoid duplication of work between the side file and the index builder, side
+file updates are elided when the record ID for the update is greater than the
+cursor position within the record ID space.
+
+To minimize changes to the rest of the codebase, XFS online repair keeps the
+replacement index hidden until it's completely ready to go.
+In other words, there is no attempt to expose the keyspace of the new index
+while repair is running.
+The complexity of such an approach would be very high and perhaps more
+appropriate to building *new* indices.
+
+**Question**: Can the full scan and live update code used to facilitate a
+repair also be used to implement a comprehensive check?
+
+*Answer*: Probably, though this has not been yet been studied.
+
+Summary Information
+```````````````````
+
+Metadata structures in this last category summarize the contents of primary
+metadata records.
+These are often used to speed up resource usage queries, and are many times
+smaller than the primary metadata which they represent.
+Check and repair both require full filesystem scans, but resource and lock
+acquisition follow the same paths as regular filesystem accesses.
+
+The superblock summary counters have special requirements due to the underlying
+implementation of the incore counters, and will be treated separately.
+Check and repair of the other types of summary counters (quota resource counts
+and file link counts) employ the same filesystem scanning and hooking
+techniques as outlined above, but because the underlying data are sets of
+integer counters, the staging data need not be a fully functional mirror of the
+ondisk structure.
+
+Inspiration for quota and file link count repair strategies were drawn from
+sections 2.12 ("Online Index Operations") through 2.14 ("Incremental View
+Maintenace") of G.  Graefe, `"Concurrent Queries and Updates in Summary Views
+and Their Indexes"
+<http://www.odbms.org/wp-content/uploads/2014/06/Increment-locks.pdf>`_, 2011.
+
+Since quotas are non-negative integer counts of resource usage, online
+quotacheck can use the incremental view deltas described in section 2.14 to
+track pending changes to the block and inode usage counts in each transaction,
+and commit those changes to a dquot side file when the transaction commits.
+Delta tracking is necessary for dquots because the index builder scans inodes,
+whereas the data structure being rebuilt is an index of dquots.
+Link count checking combines the view deltas and commit step into one because
+it sets attributes of the objects being scanned instead of writing them to a
+separate data structure.
+Each online fsck function will be discussed as case studies later in this
+document.
+
+Risk Management
+---------------
+
+During the development of online fsck, several risk factors were identified
+that may make the feature unsuitable for certain distributors and users.
+Steps can be taken to mitigate or eliminate those risks, though at a cost to
+functionality.
+
+- **Decreased performance**: Adding metadata indices to the filesystem
+  increases the time cost of persisting changes to disk, and the reverse space
+  mapping and directory parent pointers are no exception.
+  System administrators who require the maximum performance can disable the
+  reverse mapping features at format time, though this choice dramatically
+  reduces the ability of online fsck to find inconsistencies and repair them.
+
+- **Incorrect repairs**: As with all software, there might be defects in the
+  software that result in incorrect repairs being written to the filesystem.
+  Systematic fuzz testing (detailed in the next section) is employed by the
+  authors to find bugs early, but it might not catch everything.
+  The kernel build system provides Kconfig options (``CONFIG_XFS_ONLINE_SCRUB``
+  and ``CONFIG_XFS_ONLINE_REPAIR``) to enable distributors to choose not to
+  accept this risk.
+  The xfsprogs build system has a configure option (``--enable-scrub=no``) that
+  disables building of the ``xfs_scrub`` binary, though this is not a risk
+  mitigation if the kernel functionality remains enabled.
+
+- **Inability to repair**: Sometimes, a filesystem is too badly damaged to be
+  repairable.
+  If the keyspaces of several metadata indices overlap in some manner but a
+  coherent narrative cannot be formed from records collected, then the repair
+  fails.
+  To reduce the chance that a repair will fail with a dirty transaction and
+  render the filesystem unusable, the online repair functions have been
+  designed to stage and validate all new records before committing the new
+  structure.
+
+- **Misbehavior**: Online fsck requires many privileges -- raw IO to block
+  devices, opening files by handle, ignoring Unix discretionary access control,
+  and the ability to perform administrative changes.
+  Running this automatically in the background scares people, so the systemd
+  background service is configured to run with only the privileges required.
+  Obviously, this cannot address certain problems like the kernel crashing or
+  deadlocking, but it should be sufficient to prevent the scrub process from
+  escaping and reconfiguring the system.
+  The cron job does not have this protection.
+
+- **Fuzz Kiddiez**: There are many people now who seem to think that running
+  automated fuzz testing of ondisk artifacts to find mischevious behavior and
+  spraying exploit code onto the public mailing list for instant zero-day
+  disclosure is somehow of some social benefit.
+  In the view of this author, the benefit is realized only when the fuzz
+  operators help to **fix** the flaws, but this opinion apparently is not
+  widely shared among security "researchers".
+  The XFS maintainers' continuing ability to manage these events presents an
+  ongoing risk to the stability of the development process.
+  Automated testing should front-load some of the risk while the feature is
+  considered EXPERIMENTAL.
+
+Many of these risks are inherent to software programming.
+Despite this, it is hoped that this new functionality will prove useful in
+reducing unexpected downtime.

[PATCH 03/14] xfs: document the testing plan for online fsck

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Start the third chapter of the online fsck design documentation.  This
covers the testing plan to make sure that both online and offline fsck
can detect arbitrary problems and correct them without making things
worse.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  187 ++++++++++++++++++++
 1 file changed, 187 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index a03a7b9f0250..d630b6bdbe4a 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -563,3 +563,190 @@ functionality.
 Many of these risks are inherent to software programming.
 Despite this, it is hoped that this new functionality will prove useful in
 reducing unexpected downtime.
+
+3. Testing Plan
+===============
+
+As stated before, fsck tools have three main goals:
+
+1. Detect inconsistencies in the metadata;
+
+2. Eliminate those inconsistencies; and
+
+3. Minimize further loss of data.
+
+Demonstrations of correct operation are necessary to build users' confidence
+that the software behaves within expectations.
+Unfortunately, it was not really feasible to perform regular exhaustive testing
+of every aspect of a fsck tool until the introduction of low-cost virtual
+machines with high-IOPS storage.
+With ample hardware availability in mind, the testing strategy for the online
+fsck project involves differential analysis against the existing fsck tools and
+systematic testing of every attribute of every type of metadata object.
+Testing can be split into four major categories, as discussed below.
+
+Integrated Testing with fstests
+-------------------------------
+
+The primary goal of any free software QA effort is to make testing as
+inexpensive and widespread as possible to maximize the scaling advantages of
+community.
+In other words, testing should maximize the breadth of filesystem configuration
+scenarios and hardware setups.
+This improves code quality by enabling the authors of online fsck to find and
+fix bugs early, and helps developers of new features to find integration
+issues earlier in their development effort.
+
+The Linux filesystem community shares a common QA testing suite,
+`fstests <https://git.kernel.org/pub/scm/fs/xfs/xfstests-dev.git/>`_, for
+functional and regression testing.
+Even before development work began on online fsck, fstests (when run on XFS)
+would run both the ``xfs_check`` and ``xfs_repair -n`` commands on the test and
+scratch filesystems between each test.
+This provides a level of assurance that the kernel and the fsck tools stay in
+alignment about what constitutes consistent metadata.
+During development of the online checking code, fstests was modified to run
+``xfs_scrub -n`` between each test to ensure that the new checking code
+produces the same results as the two existing fsck tools.
+
+To start development of online repair, fstests was modified to run
+``xfs_repair`` to rebuild the filesystem's metadata indices between tests.
+This ensures that offline repair does not crash, leave a corrupt filesystem
+after it exists, or trigger complaints from the online check.
+This also established a baseline for what can and cannot be repaired offline.
+To complete the first phase of development of online repair, fstests was
+modified to be able to run ``xfs_scrub`` in a "force rebuild" mode.
+This enables a comparison of the effectiveness of online repair as compared to
+the existing offline repair tools.
+
+General Fuzz Testing of Metadata Blocks
+---------------------------------------
+
+XFS benefits greatly from having a very robust debugging tool, ``xfs_db``.
+
+Before development of online fsck even began, a set of fstests were created
+to test the rather common fault that entire metadata blocks get corrupted.
+This required the creation of fstests library code that can create a filesystem
+containing every possible type of metadata object.
+Next, individual test cases were created to create a test filesystem, identify
+a single block of a specific type of metadata object, trash it with the
+existing ``blocktrash`` command in ``xfs_db``, and test the reaction of a
+particular metadata validation strategy.
+
+This earlier test suite enabled XFS developers to test the ability of the
+in-kernel validation functions and the ability of the offline fsck tool to
+detect and eliminate the inconsistent metadata.
+This part of the test suite was extended to cover online fsck in exactly the
+same manner.
+
+In other words, for a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem:
+
+  * Write garbage to it
+
+  * Test the reactions of:
+
+    1. The kernel verifiers to stop obviously bad metadata
+    2. Offline repair (``xfs_repair``) to detect and fix
+    3. Online repair (``xfs_scrub``) to detect and fix
+
+Targeted Fuzz Testing of Metadata Records
+-----------------------------------------
+
+A quick conversation with the other XFS developers revealed that the existing
+test infrastructure could be extended to provide a much more powerful
+facility: targeted fuzz testing of every metadata field of every metadata
+object in the filesystem.
+``xfs_db`` can modify every field of every metadata structure in every
+block in the filesystem to simulate the effects of memory corruption and
+software bugs.
+Given that fstests already contains the ability to create a filesystem
+containing every metadata format known to the filesystem, ``xfs_db`` can be
+used to perform exhaustive fuzz testing!
+
+For a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem...
+
+  * For each record inside that metadata object...
+
+    * For each field inside that record...
+
+      * For each conceivable type of transformation that can be applied to a bit field...
+
+        1. Clear all bits
+        2. Set all bits
+        3. Toggle the most significant bit
+        4. Toggle the middle bit
+        5. Toggle the least significant bit
+        6. Add a small quantity
+        7. Subtract a small quantity
+        8. Randomize the contents
+
+        * ...test the reactions of:
+
+          1. The kernel verifiers to stop obviously bad metadata
+          2. Offline checking (``xfs_repair -n``)
+          3. Offline repair (``xfs_repair``)
+          4. Online checking (``xfs_scrub -n``)
+          5. Online repair (``xfs_scrub``)
+          6. Both repair tools (``xfs_scrub`` and then ``xfs_repair`` if online repair doesn't succeed)
+
+This is quite the combinatoric explosion!
+
+Fortunately, having this much test coverage makes it easy for XFS developers to
+check the responses of XFS' fsck tools.
+Since the introduction of the fuzz testing framework, these tests have been
+used to discover incorrect repair code and missing functionality for entire
+classes of metadata objects in ``xfs_repair``.
+The enhanced testing was used to finalize the deprecation of ``xfs_check`` by
+confirming that ``xfs_repair`` could detect at least as many corruptions as
+the older tool.
+
+These tests have been very valuable for ``xfs_scrub`` in the same ways -- they
+allow the online fsck developers to compare online fsck against offline fsck,
+and they enable XFS developers to find deficiencies in the code base.
+
+Proposed patchsets include
+`general fuzzer improvements
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=fuzzer-improvements>`_,
+`fuzzing baselines
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=fuzz-baseline>`_,
+and `improvements in fuzz testing comprehensiveness
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=more-fuzz-testing>`_.
+
+Stress Testing
+--------------
+
+A unique requirement to online fsck is the ability to operate on a filesystem
+concurrently with regular workloads.
+Although it is of course impossible to run ``xfs_scrub`` with *zero* observable
+impact on the running system, the online repair code should never introduce
+inconsistencies into the filesystem metadata, and regular workloads should
+never notice resource starvation.
+To verify that these conditions are being met, fstests has been enhanced in
+the following ways:
+
+* For each scrub item type, create a test to exercise checking that item type
+  while running ``fsstress``.
+* For each scrub item type, create a test to exercise repairing that item type
+  while running ``fsstress``.
+* Race ``fsstress`` and ``xfs_scrub -n`` to ensure that checking the whole
+  filesystem doesn't cause problems.
+* Race ``fsstress`` and ``xfs_scrub`` in force-rebuild mode to ensure that
+  force-repairing the whole filesystem doesn't cause problems.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  freezing and thawing the filesystem.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  remounting the filesystem read-only and read-write.
+* The same, but running ``fsx`` instead of ``fsstress``.  (Not done yet?)
+
+Success is defined by the ability to run all of these tests without observing
+any unexpected filesystem shutdowns due to corrupted metadata, kernel hang
+check warnings, or any other sort of mischief.
+
+Proposed patchsets include `general stress testing
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=race-scrub-and-mount-state-changes>`_
+and the `evolution of existing per-function stress testing
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=refactor-scrub-stress>`_.

Re: [PATCH 03/14] xfs: document the testing plan for online fsck

[Dave Chinner]

On Sun, Aug 07, 2022 at 11:30:22AM -0700, Darrick J. Wong wrote:
> From: Darrick J. Wong <djwong@kernel.org>
> 
> Start the third chapter of the online fsck design documentation.  This
> covers the testing plan to make sure that both online and offline fsck
> can detect arbitrary problems and correct them without making things
> worse.
> 
> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
> ---
>  .../filesystems/xfs-online-fsck-design.rst         |  187 ++++++++++++++++++++
>  1 file changed, 187 insertions(+)


....
> +Stress Testing
> +--------------
> +
> +A unique requirement to online fsck is the ability to operate on a filesystem
> +concurrently with regular workloads.
> +Although it is of course impossible to run ``xfs_scrub`` with *zero* observable
> +impact on the running system, the online repair code should never introduce
> +inconsistencies into the filesystem metadata, and regular workloads should
> +never notice resource starvation.
> +To verify that these conditions are being met, fstests has been enhanced in
> +the following ways:
> +
> +* For each scrub item type, create a test to exercise checking that item type
> +  while running ``fsstress``.
> +* For each scrub item type, create a test to exercise repairing that item type
> +  while running ``fsstress``.
> +* Race ``fsstress`` and ``xfs_scrub -n`` to ensure that checking the whole
> +  filesystem doesn't cause problems.
> +* Race ``fsstress`` and ``xfs_scrub`` in force-rebuild mode to ensure that
> +  force-repairing the whole filesystem doesn't cause problems.
> +* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
> +  freezing and thawing the filesystem.
> +* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
> +  remounting the filesystem read-only and read-write.
> +* The same, but running ``fsx`` instead of ``fsstress``.  (Not done yet?)

I had a thought when reading this that we want to ensure that online
repair handles concurrent grow/shrink operations so that doesn't
cause problems, as well as dealing with concurrent attempts to run
independent online repair processes.

Not sure that comes under stress testing, but it was the "test while
freeze/thaw" that triggered me to think of this, so that's where I'm
commenting about it. :)

Cheers,

Dave.
-- 
Dave Chinner
david@fromorbit.com

Re: [PATCH 03/14] xfs: document the testing plan for online fsck

[Darrick J. Wong]

On Thu, Aug 11, 2022 at 10:09:45AM +1000, Dave Chinner wrote:
> On Sun, Aug 07, 2022 at 11:30:22AM -0700, Darrick J. Wong wrote:
> > From: Darrick J. Wong <djwong@kernel.org>
> > 
> > Start the third chapter of the online fsck design documentation.  This
> > covers the testing plan to make sure that both online and offline fsck
> > can detect arbitrary problems and correct them without making things
> > worse.
> > 
> > Signed-off-by: Darrick J. Wong <djwong@kernel.org>
> > ---
> >  .../filesystems/xfs-online-fsck-design.rst         |  187 ++++++++++++++++++++
> >  1 file changed, 187 insertions(+)
> 
> 
> ....
> > +Stress Testing
> > +--------------
> > +
> > +A unique requirement to online fsck is the ability to operate on a filesystem
> > +concurrently with regular workloads.
> > +Although it is of course impossible to run ``xfs_scrub`` with *zero* observable
> > +impact on the running system, the online repair code should never introduce
> > +inconsistencies into the filesystem metadata, and regular workloads should
> > +never notice resource starvation.
> > +To verify that these conditions are being met, fstests has been enhanced in
> > +the following ways:
> > +
> > +* For each scrub item type, create a test to exercise checking that item type
> > +  while running ``fsstress``.
> > +* For each scrub item type, create a test to exercise repairing that item type
> > +  while running ``fsstress``.
> > +* Race ``fsstress`` and ``xfs_scrub -n`` to ensure that checking the whole
> > +  filesystem doesn't cause problems.
> > +* Race ``fsstress`` and ``xfs_scrub`` in force-rebuild mode to ensure that
> > +  force-repairing the whole filesystem doesn't cause problems.
> > +* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
> > +  freezing and thawing the filesystem.
> > +* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
> > +  remounting the filesystem read-only and read-write.
> > +* The same, but running ``fsx`` instead of ``fsstress``.  (Not done yet?)
> 
> I had a thought when reading this that we want to ensure that online
> repair handles concurrent grow/shrink operations so that doesn't
> cause problems, as well as dealing with concurrent attempts to run
> independent online repair processes.
> 
> Not sure that comes under stress testing, but it was the "test while
> freeze/thaw" that triggered me to think of this, so that's where I'm
> commenting about it. :)

Hmm.  I hadn't really given that much thought.  Let me go add that to
the test suite and see how many daemons come pouring out...

--D

> Cheers,
> 
> Dave.
> -- 
> Dave Chinner
> david@fromorbit.com

[PATCH 04/14] xfs: document the user interface for online fsck

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Start the fourth chapter of the online fsck design documentation, which
discusses the user interface and the background scrubbing service.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  114 ++++++++++++++++++++
 1 file changed, 114 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index d630b6bdbe4a..42e82971e036 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -750,3 +750,117 @@ Proposed patchsets include `general stress testing
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=race-scrub-and-mount-state-changes>`_
 and the `evolution of existing per-function stress testing
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=refactor-scrub-stress>`_.
+
+4. User Interface
+=================
+
+The primary user of online fsck is the system administrator, just like offline
+repair.
+Online fsck presents two modes of operation to administrators:
+A foreground CLI process for online fsck on demand, and a background service
+that performs autonomous checking and repair.
+
+Checking on Demand
+------------------
+
+For administrators who want the absolute freshest information about the
+metadata in a filesystem, ``xfs_scrub`` can be run as a foreground process on
+a command line.
+The program checks every piece of metadata in the filesystem while the
+administrator waits for the results to be reported, just like the existing
+``xfs_repair`` tool.
+Both tools share a ``-n`` option to perform a read-only scan, and a ``-v``
+option to increase the verbosity of the information reported.
+
+A new feature of ``xfs_scrub`` is the ``-x`` option, which employs the error
+correction capabilities of the hardware to check data file contents.
+The media scan is not enabled by default because it may dramatically increase
+program runtime and consume a lot of bandwidth on older storage hardware.
+
+The output of a foreground invocation is captured in the system log.
+
+The ``xfs_scrub_all`` program walks the list of mounted filesystems and
+initiates ``xfs_scrub`` for each of them in parallel.
+It serializes scans for any filesystems that resolve to the same top level
+kernel block device to prevent resource overconsumption.
+
+Background Service
+------------------
+
+To reduce the workload of system administrators, the ``xfs_scrub`` package
+provides a suite of `systemd <https://systemd.io/>`_ timers and services that
+run online fsck automatically on weekends.
+The background service configures scrub to run with as little privilege as
+possible, the lowest CPU and IO priority, and in a CPU-constrained single
+threaded mode.
+It is hoped that this minimizes the amount of load generated on the system and
+avoids starving regular workloads.
+
+The output of the background service is also captured in the system log.
+If desired, reports of failures (either due to inconsistencies or mere runtime
+errors) can be emailed automatically by setting the ``EMAIL_ADDR`` environment
+variable in the following service files:
+
+* ``xfs_scrub_fail@.service``
+* ``xfs_scrub_media_fail@.service``
+* ``xfs_scrub_all_fail.service``
+
+The decision to enable the background scan is left to the system administrator.
+This can be done by enabling either of the following services:
+
+* ``xfs_scrub_all.timer`` on systemd systems
+* ``xfs_scrub_all.cron`` on non-systemd systems
+
+This automatic weekly scan is configured out of the box to perform an
+additional media scan of all file data once per month.
+This is less foolproof than, say, storing file data block checksums, but much
+more performant if application software provides its own integrity checking,
+redundancy can be provided elsewhere above the filesystem, or the storage
+device's integrity guarantees are deemed sufficient.
+
+The systemd unit file definitions have been subjected to a security audit
+(as of systemd 249) to ensure that the xfs_scrub processes have as little
+access to the rest of the system as possible.
+This was performed via ``systemd-analyze security``, after which privileges
+were restricted to the minimum required, sandboxing was set up to the maximal
+extent possible with sandboxing and system call filtering; and access to the
+filesystem tree was restricted to the minimum needed to start the program and
+access the filesystem being scanned.
+The service definition files restrict CPU usage to 80% of one CPU core, and
+apply as nice of a priority to IO and CPU scheduling as possible.
+This measure was taken to minimize delays in the rest of the filesystem.
+No such hardening has been performed for the cron job.
+
+Proposed patchset:
+`Enabling the xfs_scrub background service
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-media-scan-service>`_.
+
+Health Reporting
+----------------
+
+XFS caches a summary of each filesystem's health status in memory.
+The information is updated whenever ``xfs_scrub`` is run, or whenever
+inconsistencies are detected in the filesystem metadata during regular
+operations.
+System administrators should use the ``health`` command of ``xfs_spaceman`` to
+download this information into a human-readable format.
+If problems have been observed, the administrator can schedule a reduced
+service window to run the online repair tool to correct the problem.
+Failing that, the administrator can decide to schedule a maintenance window to
+run the traditional offline repair tool to correct the problem.
+
+**Question**: Should the health reporting integrate with the new inotify fs
+error notification system?
+
+**Question**: Would it be helpful for sysadmins to have a daemon to listen for
+corruption notifications and initiate a repair?
+
+*Answer*: These questions remain unanswered, but should be a part of the
+conversation with early adopters and potential downstream users of XFS.
+
+Proposed patchsets include
+`wiring up health reports to correction returns
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=corruption-health-reports>`_
+and
+`preservation of sickness info during memory reclaim
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=indirect-health-reporting>`_.

Re: [PATCH 04/14] xfs: document the user interface for online fsck

[Dave Chinner]

On Sun, Aug 07, 2022 at 11:30:28AM -0700, Darrick J. Wong wrote:
> From: Darrick J. Wong <djwong@kernel.org>
> 
> Start the fourth chapter of the online fsck design documentation, which
> discusses the user interface and the background scrubbing service.
> 
> Signed-off-by: Darrick J. Wong <djwong@kernel.org>
> ---
>  .../filesystems/xfs-online-fsck-design.rst         |  114 ++++++++++++++++++++
>  1 file changed, 114 insertions(+)
> 
> 
> diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
> index d630b6bdbe4a..42e82971e036 100644
> --- a/Documentation/filesystems/xfs-online-fsck-design.rst
> +++ b/Documentation/filesystems/xfs-online-fsck-design.rst
> @@ -750,3 +750,117 @@ Proposed patchsets include `general stress testing
>  <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=race-scrub-and-mount-state-changes>`_
>  and the `evolution of existing per-function stress testing
>  <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=refactor-scrub-stress>`_.
> +
> +4. User Interface
> +=================
> +
> +The primary user of online fsck is the system administrator, just like offline
> +repair.
> +Online fsck presents two modes of operation to administrators:
> +A foreground CLI process for online fsck on demand, and a background service
> +that performs autonomous checking and repair.
> +
> +Checking on Demand
> +------------------
> +
> +For administrators who want the absolute freshest information about the
> +metadata in a filesystem, ``xfs_scrub`` can be run as a foreground process on
> +a command line.
> +The program checks every piece of metadata in the filesystem while the
> +administrator waits for the results to be reported, just like the existing
> +``xfs_repair`` tool.
> +Both tools share a ``-n`` option to perform a read-only scan, and a ``-v``
> +option to increase the verbosity of the information reported.
> +
> +A new feature of ``xfs_scrub`` is the ``-x`` option, which employs the error
> +correction capabilities of the hardware to check data file contents.
> +The media scan is not enabled by default because it may dramatically increase
> +program runtime and consume a lot of bandwidth on older storage hardware.

So '-x' runs a media scrub command? What does that do with software
RAID? Does that trigger parity checks of the RAID volume, or pass
through to the underlying hardware to do physical media scrub? Or
maybe both?

Rewriting the paragraph to be focussed around the functionality
being provided (i.e "media scrubbing is a new feature of xfs_scrub.
It provides .....")

> +The output of a foreground invocation is captured in the system log.

At what log level?

> +The ``xfs_scrub_all`` program walks the list of mounted filesystems and
> +initiates ``xfs_scrub`` for each of them in parallel.
> +It serializes scans for any filesystems that resolve to the same top level
> +kernel block device to prevent resource overconsumption.

Is this serialisation necessary for non-HDD devices?

> +Background Service
> +------------------
> +
> +To reduce the workload of system administrators, the ``xfs_scrub`` package
> +provides a suite of `systemd <https://systemd.io/>`_ timers and services that
> +run online fsck automatically on weekends.

Weekends change depending on where you are in the world, right? So
maybe this should be more explicit?

[....]

> +**Question**: Should the health reporting integrate with the new inotify fs
> +error notification system?

Can the new inotify fs error notification system report complex
health information structures? How much pain is involved in making
it do what we want, considering we already have a health reporting
ioctl that can be polled?

> +**Question**: Would it be helpful for sysadmins to have a daemon to listen for
> +corruption notifications and initiate a repair?

Seems like an obvious extension to the online repair capability.

Cheers,

Dave.
-- 
Dave Chinner
david@fromorbit.com

Re: [PATCH 04/14] xfs: document the user interface for online fsck

[Darrick J. Wong]

On Thu, Aug 11, 2022 at 10:20:12AM +1000, Dave Chinner wrote:
> On Sun, Aug 07, 2022 at 11:30:28AM -0700, Darrick J. Wong wrote:
> > From: Darrick J. Wong <djwong@kernel.org>
> > 
> > Start the fourth chapter of the online fsck design documentation, which
> > discusses the user interface and the background scrubbing service.
> > 
> > Signed-off-by: Darrick J. Wong <djwong@kernel.org>
> > ---
> >  .../filesystems/xfs-online-fsck-design.rst         |  114 ++++++++++++++++++++
> >  1 file changed, 114 insertions(+)
> > 
> > 
> > diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
> > index d630b6bdbe4a..42e82971e036 100644
> > --- a/Documentation/filesystems/xfs-online-fsck-design.rst
> > +++ b/Documentation/filesystems/xfs-online-fsck-design.rst
> > @@ -750,3 +750,117 @@ Proposed patchsets include `general stress testing
> >  <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=race-scrub-and-mount-state-changes>`_
> >  and the `evolution of existing per-function stress testing
> >  <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=refactor-scrub-stress>`_.
> > +
> > +4. User Interface
> > +=================
> > +
> > +The primary user of online fsck is the system administrator, just like offline
> > +repair.
> > +Online fsck presents two modes of operation to administrators:
> > +A foreground CLI process for online fsck on demand, and a background service
> > +that performs autonomous checking and repair.
> > +
> > +Checking on Demand
> > +------------------
> > +
> > +For administrators who want the absolute freshest information about the
> > +metadata in a filesystem, ``xfs_scrub`` can be run as a foreground process on
> > +a command line.
> > +The program checks every piece of metadata in the filesystem while the
> > +administrator waits for the results to be reported, just like the existing
> > +``xfs_repair`` tool.
> > +Both tools share a ``-n`` option to perform a read-only scan, and a ``-v``
> > +option to increase the verbosity of the information reported.
> > +
> > +A new feature of ``xfs_scrub`` is the ``-x`` option, which employs the error
> > +correction capabilities of the hardware to check data file contents.
> > +The media scan is not enabled by default because it may dramatically increase
> > +program runtime and consume a lot of bandwidth on older storage hardware.
> 
> So '-x' runs a media scrub command? What does that do with software
> RAID?

Nothing special unless the RAID controller itself does parity checking
of reads -- the kernel doesn't have any API calls (that I know of) to do
that.  I think md-raid5 will check the parity, but afaict nothing else
(raid1) does that.

> Does that trigger parity checks of the RAID volume, or pass
> through to the underlying hardware to do physical media scrub?

Chaitanya proposed a userspace api so that xfs_scrub could actually ask
the hardware to perform a media verification[1], but willy pointed out
that it none of the device protocols have a means for the device to
prove that it did anything, so it stalled.

[1] https://lore.kernel.org/linux-fsdevel/20220713072019.5885-1-kch@nvidia.com/

> Or maybe both?

I wish. :)

> Rewriting the paragraph to be focussed around the functionality
> being provided (i.e "media scrubbing is a new feature of xfs_scrub.
> It provides .....")

Er.. you're doing that, or asking me to do it?

> > +The output of a foreground invocation is captured in the system log.
> 
> At what log level?

That depends on the message, but right now it only uses
LOG_{ERR,WARNING,INFO}.

Errors, corruptions, and unfixable problems are LOG_ERR.

Warnings are LOG_WARNING.

Notices of infomration, repairs completed, and optimizations made are
all recorded with LOG_INFO.

> > +The ``xfs_scrub_all`` program walks the list of mounted filesystems and
> > +initiates ``xfs_scrub`` for each of them in parallel.
> > +It serializes scans for any filesystems that resolve to the same top level
> > +kernel block device to prevent resource overconsumption.
> 
> Is this serialisation necessary for non-HDD devices?

That ultimately depends on the preferences of the sysadmins, but for the
initial push I'd rather err on the side of using fewer iops on a running
system.

> > +Background Service
> > +------------------
> > +
> > +To reduce the workload of system administrators, the ``xfs_scrub`` package
> > +provides a suite of `systemd <https://systemd.io/>`_ timers and services that
> > +run online fsck automatically on weekends.
> 
> Weekends change depending on where you are in the world, right? So
> maybe this should be more explicit?

Sunday at 3:10am, whenever that is in the local time zone.

> [....]
> 
> > +**Question**: Should the health reporting integrate with the new inotify fs
> > +error notification system?
> 
> Can the new inotify fs error notification system report complex
> health information structures?

In theory, yes, said the authors.

> How much pain is involved in making
> it do what we want, considering we already have a health reporting
> ioctl that can be polled?

I haven't tried this myself, but I think it involves defining a new type
code and message length within the inotify system.  The last time I
looked at the netlink protocol, I /think/ I saw that it's the case that
the consuming programs will read the header, see that there's a type
code and a buffer length, and decide to use it or skip it.

That said, there were some size and GFP_ limits on what could be sent,
so I don't know how difficult it would be to make this part actually
work in practice.  Gabriel said it wouldn't be too difficult once I was
ready.

> > +**Question**: Would it be helpful for sysadmins to have a daemon to listen for
> > +corruption notifications and initiate a repair?
> 
> Seems like an obvious extension to the online repair capability.

...too bad there are dragons thataways.

--D

> Cheers,
> 
> Dave.
> -- 
> Dave Chinner
> david@fromorbit.com

[PATCH 05/14] xfs: document the filesystem metadata checking strategy

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Begin the fifth chapter of the online fsck design documentation, where
we discuss the details of the data structures and algorithms used by the
kernel to examine filesystem metadata and cross-reference it around the
filesystem.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  579 ++++++++++++++++++++
 .../filesystems/xfs-self-describing-metadata.rst   |    1 
 2 files changed, 580 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 42e82971e036..f45bf97fa9c4 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -864,3 +864,582 @@ Proposed patchsets include
 and
 `preservation of sickness info during memory reclaim
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=indirect-health-reporting>`_.
+
+5. Kernel Algorithms and Data Structures
+========================================
+
+This section discusses the key algorithms and data structures of the kernel
+code that provide the ability to check and repair metadata while the system
+is running.
+The first chapters in this section reveal the pieces that provide the
+foundation for checking metadata.
+The remainder of this section presents the mechanisms through which XFS
+regenerates itself.
+
+Self Describing Metadata
+------------------------
+
+Starting with XFS version 5 in 2012, XFS updated the format of nearly every
+ondisk block header to record a magic number, a checksum, a universally
+"unique" identifier (UUID), an owner code, the ondisk address of the block,
+and a log sequence number.
+When loading a block buffer from disk, the magic number, UUID, owner, and
+ondisk address confirm that the retrieved block matches the specific owner of
+the current filesystem, and that the information contained in the block is
+supposed to be found at the ondisk address.
+The first three components enable checking tools to disregard alleged metadata
+that doesn't belong to the filesystem, and the fourth component enables the
+filesystem to detect lost writes.
+
+The logging code maintains the checksum and the log sequence number of the last
+transactional update.
+Checksums are useful for detecting torn writes and other mischief between the
+computer and its storage devices.
+Sequence number tracking enables log recovery to avoid applying out of date
+log updates to the filesystem.
+
+These two features improve overall runtime resiliency by providing a means for
+the filesystem to detect obvious corruption when reading metadata blocks from
+disk, but these buffer verifiers cannot provide any consistency checking
+between metadata structures.
+
+For more information, please see the documentation for
+Documentation/filesystems/xfs-self-describing-metadata.rst
+
+Reverse Mapping
+---------------
+
+The original design of XFS (circa 1993) is an improvement upon 1980s Unix
+filesystem design.
+In those days, storage density was expensive, CPU time was scarce, and
+excessive seek time could kill performance.
+For performance reasons, filesystem authors were reluctant to add redundancy to
+the filesystem, even at the cost of data integrity.
+Filesystems designers in the early 21st century choose different strategies to
+increase internal redundancy -- either storing nearly identical copies of
+metadata, or more space-efficient techniques such as erasure coding.
+Obvious corruptions are typically repaired by copying replicas or
+reconstructing from codes.
+
+For XFS, a different redundancy strategy was chosen to modernize the design:
+a secondary space usage index that maps allocated disk extents back to their
+owners.
+By adding a new index, the filesystem retains most of its ability to scale
+well to heavily threaded workloads involving large datasets, since the primary
+file metadata (the directory tree, the file block map, and the allocation
+groups) remain unchanged.
+Although the reverse-mapping feature increases overhead costs for space
+mapping activities just like any other system that improves redundancy, it
+has two critical advantages: first, the reverse index is key to enabling online
+fsck and other requested functionality such as filesystem reorganization,
+better media failure reporting, and shrinking.
+Second, the different ondisk storage format of the reverse mapping btree
+defeats device-level deduplication, because the filesystem requires real
+redundancy.
+
+A criticism of adding the secondary index is that it does nothing to improve
+the robustness of user data storage itself.
+This is a valid point, but adding a new index for file data block checksums
+increases write amplification and turns data overwrites into copy-writes, which
+age the filesystem prematurely.
+In keeping with thirty years of precedent, users who want file data integrity
+can supply as powerful a solution as they require.
+As for metadata, the complexity of adding a new secondary index of space usage
+is much less than adding volume management and storage device mirroring to XFS
+itself.
+Perfection of RAID and volume management are best left to existing layers in
+the kernel.
+
+The information captured in a reverse space mapping record is as follows:
+
+.. code-block:: c
+
+	struct xfs_rmap_irec {
+	    xfs_agblock_t    rm_startblock;   /* extent start block */
+	    xfs_extlen_t     rm_blockcount;   /* extent length */
+	    uint64_t         rm_owner;        /* extent owner */
+	    uint64_t         rm_offset;       /* offset within the owner */
+	    unsigned int     rm_flags;        /* state flags */
+	};
+
+The first two fields capture the location and size of the physical space,
+in units of filesystem blocks.
+The owner field tells scrub which metadata structure or file inode have been
+assigned this space.
+For space allocated to files, the offset field tells scrub where the space was
+mapped within the file fork.
+Finally, the flags field provides extra information about the space usage --
+is this an attribute fork extent?  A file mapping btree extent?  Or an
+unwritten data extent?
+
+Online filesystem checking judges the consistency of each primary metadata
+record by comparing its information against all other space indices.
+The reverse mapping index plays a key role in the consistency checking process
+because it contains a centralized alternate copy of all space allocation
+information.
+Program runtime and ease of resource acquisition are the only real limits to
+what online checking can consult.
+For example, a file data extent mapping can be checked against:
+
+* The absence of an entry in the free space information.
+* The absence of an entry in the inode index.
+* The absence of an entry in the reference count data if the file is not
+  marked as having shared extents.
+* The correspondence of an entry in the reverse mapping information.
+
+A key observation here is that only the reverse mapping can provide a positive
+affirmation of correctness if the primary metadata is in doubt.
+The checking code for most primary metadata follows a path similar to the
+one outlined above.
+
+A second observation to make about this secondary index is that proving its
+consistency with the primary metadata is difficult.
+Demonstrating that a given reverse mapping record exactly corresponds to the
+primary space metadata involves a full scan of all primary space metadata,
+which is very time intensive.
+Scanning activity for online fsck can only use non-blocking lock acquisition
+primitives if the locking order is not the regular order as used by the rest of
+the filesystem.
+This means that forward progress during this part of a scan of the reverse
+mapping data cannot be guaranteed if system load is especially heavy.
+Therefore, it is not practical for online check to detect reverse mapping
+records that lack a counterpart in the primary metadata.
+Instead, scrub relies on rigorous cross-referencing during the primary space
+mapping structure checks.
+
+Reverse mappings also play a key role in reconstruction of primary metadata.
+The secondary information is general enough for online repair to synthesize a
+complete copy of any primary space management metadata by locking that
+resource, querying all reverse mapping indices looking for records matching
+the relevant resource, and transforming the mapping into an appropriate format.
+The details of how these records are staged, written to disk, and committed
+into the filesystem are covered in subsequent sections.
+
+Checking and Cross-Referencing
+------------------------------
+
+The first step of checking a metadata structure is to examine every record
+contained within the structure and its relationship with the rest of the
+system.
+XFS contains multiple layers of checking to try to prevent inconsistent
+metadata from wreaking havoc on the system.
+Each of these layers contributes information that helps the kernel to make
+three decisions about the health of a metadata structure:
+
+- Is a part of this structure obviously corrupt (``XFS_SCRUB_OFLAG_CORRUPT``) ?
+- Is this structure inconsistent with the rest of the system
+  (``XFS_SCRUB_OFLAG_XCORRUPT``) ?
+- Is there so much damage around the filesystem that cross-referencing is not
+  possible (``XFS_SCRUB_OFLAG_XFAIL``) ?
+- Can the structure be optimized to improve performance or reduce the size of
+  metadata (``XFS_SCRUB_OFLAG_PREEN``) ?
+- Does the structure contain data that is not inconsistent but deserves review
+  by the system administrator (``XFS_SCRUB_OFLAG_WARNING``) ?
+
+The following sections describe how the metadata scrubbing process works.
+
+Metadata Buffer Verification
+````````````````````````````
+
+The lowest layer of metadata protection in XFS are the metadata verifiers built
+into the buffer cache.
+These functions perform inexpensive internal consistency checking of the block
+itself, and answer these questions:
+
+- Does the block belong to this filesystem?
+
+- Does the block belong to the structure that asked for the read?
+  This assumes that metadata blocks only have one owner, which is always true
+  in XFS.
+
+- Is the type of data stored in the block within a reasonable range of what
+  scrub is expecting?
+
+- Does the physical location of the block match the location it was read from?
+
+- Does the block checksum match the data?
+
+The scope of the protections here are very limited -- verifiers can only
+establish that the filesystem code is reasonably free of gross corruption bugs
+and that the storage system is reasonably competent at retrieval.
+Corruption problems observed at runtime cause the generation of health reports,
+failed system calls, and in the extreme case, filesystem shutdowns if the
+corrupt metadata force the cancellation of a dirty transaction.
+
+Every online fsck scrubbing function is expected to read every ondisk metadata
+block of a structure in the course of checking the structure.
+Corruption problems observed during a check are immediately reported to
+userspace as corruption; during a cross-reference, they are reported as a
+failure to cross-reference once the full examination is complete.
+Reads satisfied by a buffer already in cache (and hence already verified)
+bypass these checks.
+
+Internal Consistency Checks
+```````````````````````````
+
+The next higher level of metadata protection is the internal record
+verification code built into the filesystem.
+These checks are split between the buffer verifiers, the in-filesystem users of
+the buffer cache, and the scrub code itself, depending on the amount of higher
+level context required.
+The scope of checking is still internal to the block.
+For performance reasons, regular code may skip some of these checks unless
+debugging is enabled or a write is about to occur.
+Scrub functions, of course, must check all possible problems.
+Either way, these higher level checking functions answer these questions:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- If the block contains records, do the records fit within the block?
+
+- If the block tracks internal free space information, is it consistent with
+  the record areas?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+Record checks in this category are more rigorous and more time-intensive.
+For example, block pointers and inumbers are checked to ensure that they point
+within the dynamically allocated parts of an allocation group and within
+the filesystem.
+Names are checked for invalid characters, and flags are checked for invalid
+combinations.
+Other record attributes are checked for sensible values.
+Btree records spanning an interval of the btree keyspace are checked for
+correct order and lack of mergeability (except for file fork mappings).
+
+Validation of Userspace-Controlled Record Attributes
+````````````````````````````````````````````````````
+
+Various pieces of filesystem metadata are directly controlled by userspace.
+Because of this nature, validation work cannot be more precise than checking
+that a value is within the possible range.
+These fields include:
+
+- Superblock fields controlled by mount options
+- Filesystem labels
+- File timestamps
+- File permissions
+- File size
+- File flags
+- Names present in directory entries, extended attribute keys, and filesystem
+  labels
+- Extended attribute key namespaces
+- Extended attribute values
+- File data block contents
+- Quota limits
+- Quota timer expiration (if resource usage exceeds the soft limit)
+
+Cross-Referencing Space Metadata
+````````````````````````````````
+
+The next higher level of checking is cross-referencing records between metadata
+structures.
+For regular runtime code, the cost of these checks is considered to be
+prohibitively expensive, but as scrub is dedicated to rooting out
+inconsistencies, it must pursue all avenues of inquiry.
+The exact set of cross-referencing is highly dependent on the context of the
+data structure being checked.
+
+The XFS btree code has keyspace scanning functions that online fsck uses to
+cross reference one structure with another.
+Specifically, scrub can scan the key space of an index to determine if that
+keyspace is fully, sparsely, or not at all mapped to records.
+For the reverse mapping btree, it is possible to mask parts of the key for the
+purposes of performing a keyspace scan so that scrub can decide if the rmap
+btree contains records mapping a certain extent of physical space without the
+sparsenses of the rest of the rmap keyspace getting in the way.
+
+Btree blocks undergo the following checks before cross-referencing:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the btree point to valid block addresses for the type
+  of btree?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each node block record, does the record key accurate reflect the contents
+  of the child block?
+
+Space allocation records are cross-referenced as follows:
+
+1. Any space mentioned by any metadata structure are cross-referenced as
+   follows:
+
+   - Does the reverse mapping index list only the appropriate owner as the
+     owner of each block?
+
+   - Are none of the blocks claimed as free space?
+
+   - If these aren't file data blocks, are none of the blocks claimed as space
+     shared by different owners?
+
+2. Btree blocks are cross-referenced as follows:
+
+   - Everything in class 1 above.
+
+   - If there's a parent node block, do the keys listed for this block match the
+     keyspace of this block?
+
+   - Do the sibling pointers point to valid blocks?  Of the same level?
+
+   - Do the child pointers point to valid blocks?  Of the next level down?
+
+3. Free space btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Does the reverse mapping index list no owners of this space?
+
+   - Is this space not claimed by the inode index for inodes?
+
+   - Is it not mentioned by the reference count index?
+
+   - Is there a matching record in the other free space btree?
+
+4. Inode btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is there a matching record in free inode btree?
+
+   - Do cleared bits in the holemask correspond with inode clusters?
+
+   - Do set bits in the freemask correspond with inode records with zero link
+     count?
+
+5. Inode records are cross-referenced as follows:
+
+   - Everything in class 1.
+
+   - Do all the fields that summarize information about the file forks actually
+     match those forks?
+
+   - Does each inode with zero link count correspond to a record in the free
+     inode btree?
+
+6. File fork space mapping records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is this space not mentioned by the inode btrees?
+
+   - If this is a CoW fork mapping, does it correspond to a CoW entry in the
+     reference count btree?
+
+7. Reference count records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Within the space subkeyspace of the rmap btree (that is to say, all
+     records mapped to a particular space extent and ignoring the owner info),
+     are there the same number of reverse mapping records for each block as the
+     reference count record claims?
+
+Proposed patchsets are the series to find gaps in
+`refcount btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-refcount-gaps>`_,
+`inode btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-inobt-gaps>`_, and
+`rmap btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-rmapbt-gaps>`_ records;
+to find
+`mergeable records
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-mergeable-records>`_;
+and to
+`improve cross referencing with rmap
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-strengthen-rmap-checking>`_
+before starting a repair.
+
+Checking Extended Attributes
+````````````````````````````
+
+Extended attributes implement a key-value store that enable fragments of data
+to be attached to any file.
+Both the kernel and userspace can access the keys and values, subject to
+namespace and privilege restrictions.
+Most typically these fragments are metadata about the file -- origins, security
+contexts, user-supplied labels, indexing information, etc.
+
+Names can be as long as 255 bytes and can exist in several different
+namespaces.
+Values can be as large as 64KB.
+A file's extended attributes are stored in blocks mapped by the attr fork.
+The mappings point to leaf blocks, remote value blocks, or dabtree blocks.
+Block 0 in the attribute fork is always the top of the structure, but otherwise
+each of the three types of blocks can be found at any offset in the attr fork.
+Leaf blocks contain attribute key records that point to the name and the value.
+Names are always stored elsewhere in the same leaf block.
+Values that are less than 3/4 the size of a filesystem block are also stored
+elsewhere in the same leaf block.
+Remote value blocks contain values that are too large to fit inside a leaf.
+If the leaf information exceeds a single filesystem block, a dabtree (also
+rooted at block 0) is created to map hashes of the attribute names to leaf
+blocks in the attr fork.
+
+Checking an extended attribute structure is not so straightfoward due to the
+lack of separation between attr blocks and index blocks.
+Scrub must read each block mapped by the attr fork and ignore the non-leaf
+blocks:
+
+1. Walk the dabtree in the attr fork (if present) to ensure that there are no
+   irregularities in the blocks or dabtree mappings that do not point to
+   attr leaf blocks.
+
+2. Walk the blocks of the attr fork looking for leaf blocks.
+   For each entry inside a leaf:
+
+   a. Validate that the name does not contain invalid characters.
+
+   b. Read the attr value.
+      This performs a named lookup of the attr name to ensure the correctness
+      of the dabtree.
+      If the value is stored in a remote block, this also validates the
+      integrity of the remote value block.
+
+Checking and Cross-Referencing Directories
+``````````````````````````````````````````
+
+The filesystem directory tree is a directed acylic graph structure, with files
+constituting the nodes, and directory entries (dirents) constituting the edges.
+Directories are a special type of file containing a set of mappings from a
+255-byte sequence (name) to an inumber.
+These are called directory entries, or dirents for short.
+Each directory file must have exactly one directory pointing to the file.
+A root directory points to itself.
+Directory entries point to files of any type.
+Each non-directory file may have multiple directories point to it.
+
+In XFS, directories are implemented as a file containing up to three 32GB
+partitions.
+The first partition contains directory entry data blocks.
+Each data block contains variable-sized records associating a user-provided
+name with an inumber and, optionally, a file type.
+If the directory entry data grows beyond one block, the second partition (which
+exists as post-EOF extents) is populated with a block containing free space
+information and an index that maps hashes of the dirent names to directory data
+blocks in the first partition.
+This makes directory name lookups very fast.
+If this second partition grows beyond one block, the third partition is
+populated with a linear array of free space information for faster
+expansions.
+If the free space has been separated and the second partition grows again
+beyond one block, then a dabtree is used to map hashes of dirent names to
+directory data blocks.
+
+Checking a directory is pretty straightfoward:
+
+1. Walk the dabtree in the second partition (if present) to ensure that there
+   are no irregularities in the blocks or dabtree mappings that do not point to
+   dirent blocks.
+
+2. Walk the blocks of the first partition looking for directory entries.
+   Each dirent is checked as follows:
+
+   a. Does the name contain no invalid characters?
+
+   b. Does the inumber correspond to an actual, allocated inode?
+
+   c. Does the child inode have a nonzero link count?
+
+   d. If a file type is included in the dirent, does it match the type of the
+      inode?
+
+   e. If the child is a subdirectory, does the child's dotdot pointer point
+      back to the parent?
+
+   f. If the directory has a second partition, perform a named lookup of the
+      dirent name to ensure the correctness of the dabtree.
+
+3. Walk the free space list in the third partition (if present) to ensure that
+   the free spaces it describes are really unused.
+
+Checking operations involving :ref:`parents <dirparent>` and
+:ref:`file link counts <nlinks>` are discussed in more detail in later
+sections.
+
+Checking Directory/Attribute Btrees
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+As stated in previous sections, the directory/attribute btree (dabtree) index
+maps user-provided names to improve lookup times by avoiding linear scans.
+Internally, it maps a 32-bit hash of the name to a block offset within the
+appropriate file fork.
+
+The internal structure of a dabtree closely resembles the btrees that record
+fixed-size metadata records -- each dabtree block contains a magic number, a
+checksum, sibling pointers, a UUID, a tree level, and a log sequence number.
+The format of leaf and node records are the same -- each entry points to the
+next level down in the hierarchy, with dabtree node records pointing to dabtree
+leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere
+in the fork.
+
+Checking and cross-referencing the dabtree is very similar to what is done for
+space btrees:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the dabtree point to valid fork offsets for dabtree
+  blocks?
+
+- Do leaf pointers within the dabtree point to valid fork offsets for directory
+  or attr leaf blocks?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each dabtree node record, does the record key accurate reflect the
+  contents of the child dabtree block?
+
+- For each dabtree leaf record, does the record key accurate reflect the
+  contents of the directory or attr block?
+
+Cross-Referencing Summary Counters
+``````````````````````````````````
+
+XFS maintains three classes of summary counters: available resources, quota
+resource usage, and file link counts.
+
+In theory, the amount of available resources (data blocks, inodes, realtime
+extents) can be found by walking the entire filesystem.
+This would make for very slow reporting, so a transactional filesystem can
+maintain summaries of this information in the superblock.
+Cross-referencing these values against the filesystem metadata should be a
+simple matter of walking the free space and inode metadata in each AG and the
+realtime bitmap, but there are complications that will be discussed in
+:ref:`more detail <fscounters>` later.
+
+:ref:`Quota usage <quotacheck>` and :ref:`file link count <nlinks>`
+checking are sufficiently complicated to warrant separate sections.
+
+Post-Repair Reverification
+``````````````````````````
+
+After performing a repair, the checking code is run a second time to validate
+the new structure, and the results of the health assessment are recorded
+internally and returned to the calling process.
+This step is critical for enabling system administrator to monitor the status
+of the filesystem and the progress of any repairs.
+For developers, it is a useful means to judge the efficacy of error detection
+and correction in the online and offline checking tools.
diff --git a/Documentation/filesystems/xfs-self-describing-metadata.rst b/Documentation/filesystems/xfs-self-describing-metadata.rst
index b79dbf36dc94..a10c4ae6955e 100644
--- a/Documentation/filesystems/xfs-self-describing-metadata.rst
+++ b/Documentation/filesystems/xfs-self-describing-metadata.rst
@@ -1,4 +1,5 @@
 .. SPDX-License-Identifier: GPL-2.0
+.. _xfs_self_describing_metadata:
 
 ============================
 XFS Self Describing Metadata

Re: [PATCH 05/14] xfs: document the filesystem metadata checking strategy

[Dave Chinner]

On Sun, Aug 07, 2022 at 11:30:33AM -0700, Darrick J. Wong wrote:
> +Reverse Mapping
> +---------------
> +
> +The original design of XFS (circa 1993) is an improvement upon 1980s Unix
> +filesystem design.
> +In those days, storage density was expensive, CPU time was scarce, and
> +excessive seek time could kill performance.
> +For performance reasons, filesystem authors were reluctant to add redundancy to
> +the filesystem, even at the cost of data integrity.
> +Filesystems designers in the early 21st century choose different strategies to
> +increase internal redundancy -- either storing nearly identical copies of
> +metadata, or more space-efficient techniques such as erasure coding.
> +Obvious corruptions are typically repaired by copying replicas or
> +reconstructing from codes.
> +
> +For XFS, a different redundancy strategy was chosen to modernize the design:
> +a secondary space usage index that maps allocated disk extents back to their
> +owners.
> +By adding a new index, the filesystem retains most of its ability to scale
> +well to heavily threaded workloads involving large datasets, since the primary
> +file metadata (the directory tree, the file block map, and the allocation
> +groups) remain unchanged.
> +Although the reverse-mapping feature increases overhead costs for space
> +mapping activities just like any other system that improves redundancy, it
> +has two critical advantages: first, the reverse index is key to enabling online
> +fsck and other requested functionality such as filesystem reorganization,
> +better media failure reporting, and shrinking.
> +Second, the different ondisk storage format of the reverse mapping btree
> +defeats device-level deduplication, because the filesystem requires real
> +redundancy.

"defeats device-level deduplication" took me a bit of thought to
follow what you were trying to describe here.

I think what you mean is this:

"Second, the reverse mapping metadata uses a different on-disk
format for storage. This prevents storage device level deduplication
from removing redundant metadata the filesystem stores for recovery
purposes. Hence we will always have two physical copies of the
metadata we need for recovery in the storage device, regardless of
the technologies it implements."

> +A criticism of adding the secondary index is that it does nothing to improve
> +the robustness of user data storage itself.
> +This is a valid point, but adding a new index for file data block checksums
> +increases write amplification and turns data overwrites into copy-writes, which
> +age the filesystem prematurely.

And can come with a significant IO performance penalty.

> +Checking Extended Attributes
> +````````````````````````````
> +
> +Extended attributes implement a key-value store that enable fragments of data
> +to be attached to any file.
> +Both the kernel and userspace can access the keys and values, subject to
> +namespace and privilege restrictions.
> +Most typically these fragments are metadata about the file -- origins, security
> +contexts, user-supplied labels, indexing information, etc.
> +
> +Names can be as long as 255 bytes and can exist in several different
> +namespaces.
> +Values can be as large as 64KB.
> +A file's extended attributes are stored in blocks mapped by the attr fork.
> +The mappings point to leaf blocks, remote value blocks, or dabtree blocks.
> +Block 0 in the attribute fork is always the top of the structure, but otherwise
> +each of the three types of blocks can be found at any offset in the attr fork.
> +Leaf blocks contain attribute key records that point to the name and the value.
> +Names are always stored elsewhere in the same leaf block.
> +Values that are less than 3/4 the size of a filesystem block are also stored
> +elsewhere in the same leaf block.
> +Remote value blocks contain values that are too large to fit inside a leaf.
> +If the leaf information exceeds a single filesystem block, a dabtree (also
> +rooted at block 0) is created to map hashes of the attribute names to leaf
> +blocks in the attr fork.
> +
> +Checking an extended attribute structure is not so straightfoward due to the
> +lack of separation between attr blocks and index blocks.
> +Scrub must read each block mapped by the attr fork and ignore the non-leaf
> +blocks:
> +
> +1. Walk the dabtree in the attr fork (if present) to ensure that there are no
> +   irregularities in the blocks or dabtree mappings that do not point to
> +   attr leaf blocks.
> +
> +2. Walk the blocks of the attr fork looking for leaf blocks.
> +   For each entry inside a leaf:
> +
> +   a. Validate that the name does not contain invalid characters.
> +
> +   b. Read the attr value.
> +      This performs a named lookup of the attr name to ensure the correctness
> +      of the dabtree.
> +      If the value is stored in a remote block, this also validates the
> +      integrity of the remote value block.

I'm assuming that checking remote xattr blocks involves checking the headers
are all valid, the overall length is correct, they have valid rmap
entries, etc?

> +Checking and Cross-Referencing Directories
> +``````````````````````````````````````````
> +
> +The filesystem directory tree is a directed acylic graph structure, with files
> +constituting the nodes, and directory entries (dirents) constituting the edges.

"with hashed file names consituting the nodes" ?

> +Directories are a special type of file containing a set of mappings from a
> +255-byte sequence (name) to an inumber.
> +These are called directory entries, or dirents for short.
> +Each directory file must have exactly one directory pointing to the file.
> +A root directory points to itself.
> +Directory entries point to files of any type.

"point to inodes"

> +Each non-directory file may have multiple directories point to it.

"Each non directory inode may have multiple dirents point to it"

> +
> +In XFS, directories are implemented as a file containing up to three 32GB
> +partitions.

"implemented as offset-based file data"

> +The first partition contains directory entry data blocks.
> +Each data block contains variable-sized records associating a user-provided
> +name with an inumber and, optionally, a file type.
> +If the directory entry data grows beyond one block, the second partition (which
> +exists as post-EOF extents) is populated with a block containing free space
> +information and an index that maps hashes of the dirent names to directory data
> +blocks in the first partition.
> +This makes directory name lookups very fast.

"This is known as a "LEAF 1" format directory."

> +If this second partition grows beyond one block, the third partition is
> +populated with a linear array of free space information for faster
> +expansions.

s/for faster expansions/to speed up first fit searches when inserting
new dirents/.

"This is known as a "NODE" format directory."

> +If the free space has been separated and the second partition grows again
> +beyond one block, then a dabtree is used to map hashes of dirent names to
> +directory data blocks.

The second partition is changed to a dabtree format during NODE
format conversion - it's just the single root block form. It grows
as a btree from there by normal split/grow btree operations at that
point.  Hence I'm not sure this needs to be mentioned as a separate
step - the node form conversion should mention the second partition
transitions to the dabtree index format, and it's all obvious from
there...

> +Checking a directory is pretty straightfoward:
> +
> +1. Walk the dabtree in the second partition (if present) to ensure that there
> +   are no irregularities in the blocks or dabtree mappings that do not point to
> +   dirent blocks.
> +
> +2. Walk the blocks of the first partition looking for directory entries.
> +   Each dirent is checked as follows:
> +
> +   a. Does the name contain no invalid characters?
> +
> +   b. Does the inumber correspond to an actual, allocated inode?
> +
> +   c. Does the child inode have a nonzero link count?
> +
> +   d. If a file type is included in the dirent, does it match the type of the
> +      inode?
> +
> +   e. If the child is a subdirectory, does the child's dotdot pointer point
> +      back to the parent?
> +
> +   f. If the directory has a second partition, perform a named lookup of the
> +      dirent name to ensure the correctness of the dabtree.
> +
> +3. Walk the free space list in the third partition (if present) to ensure that
> +   the free spaces it describes are really unused.

Actaully, they should reflect the largest free space in the given
block, so this has to be checked against the bests array in the
given block.

Also, if we are walking the directory blocks, we should be checking
the empty ranges for valid tags and sizes, and that the largest 3
holes in the block are correctly ordered in the bests array, too.

> +
> +Checking operations involving :ref:`parents <dirparent>` and
> +:ref:`file link counts <nlinks>` are discussed in more detail in later
> +sections.
> +
> +Checking Directory/Attribute Btrees
> +^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
> +
> +As stated in previous sections, the directory/attribute btree (dabtree) index
> +maps user-provided names to improve lookup times by avoiding linear scans.
> +Internally, it maps a 32-bit hash of the name to a block offset within the
> +appropriate file fork.
> +
> +The internal structure of a dabtree closely resembles the btrees that record
> +fixed-size metadata records -- each dabtree block contains a magic number, a
> +checksum, sibling pointers, a UUID, a tree level, and a log sequence number.
> +The format of leaf and node records are the same -- each entry points to the
> +next level down in the hierarchy, with dabtree node records pointing to dabtree
> +leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere
> +in the fork.
> +
> +Checking and cross-referencing the dabtree is very similar to what is done for
> +space btrees:
> +
> +- Does the type of data stored in the block match what scrub is expecting?
> +
> +- Does the block belong to the owning structure that asked for the read?
> +
> +- Do the records fit within the block?
> +
> +- Are the records contained inside the block free of obvious corruptions?
> +
> +- Are the name hashes in the correct order?
> +
> +- Do node pointers within the dabtree point to valid fork offsets for dabtree
> +  blocks?
> +
> +- Do leaf pointers within the dabtree point to valid fork offsets for directory
> +  or attr leaf blocks?
> +
> +- Do child pointers point towards the leaves?
> +
> +- Do sibling pointers point across the same level?
> +
> +- For each dabtree node record, does the record key accurate reflect the
> +  contents of the child dabtree block?
> +
> +- For each dabtree leaf record, does the record key accurate reflect the
> +  contents of the directory or attr block?

This last one is checking that they point to a valid dirent and
check that the hash of the dirent name actually matches the hash
that points to it?

Cheers,

Dave.
-- 
Dave Chinner
david@fromorbit.com

Re: [PATCH 05/14] xfs: document the filesystem metadata checking strategy

[Darrick J. Wong]

On Thu, Aug 11, 2022 at 11:17:42AM +1000, Dave Chinner wrote:
> On Sun, Aug 07, 2022 at 11:30:33AM -0700, Darrick J. Wong wrote:
> > +Reverse Mapping
> > +---------------
> > +
> > +The original design of XFS (circa 1993) is an improvement upon 1980s Unix
> > +filesystem design.
> > +In those days, storage density was expensive, CPU time was scarce, and
> > +excessive seek time could kill performance.
> > +For performance reasons, filesystem authors were reluctant to add redundancy to
> > +the filesystem, even at the cost of data integrity.
> > +Filesystems designers in the early 21st century choose different strategies to
> > +increase internal redundancy -- either storing nearly identical copies of
> > +metadata, or more space-efficient techniques such as erasure coding.
> > +Obvious corruptions are typically repaired by copying replicas or
> > +reconstructing from codes.
> > +
> > +For XFS, a different redundancy strategy was chosen to modernize the design:
> > +a secondary space usage index that maps allocated disk extents back to their
> > +owners.
> > +By adding a new index, the filesystem retains most of its ability to scale
> > +well to heavily threaded workloads involving large datasets, since the primary
> > +file metadata (the directory tree, the file block map, and the allocation
> > +groups) remain unchanged.
> > +Although the reverse-mapping feature increases overhead costs for space
> > +mapping activities just like any other system that improves redundancy, it
> > +has two critical advantages: first, the reverse index is key to enabling online
> > +fsck and other requested functionality such as filesystem reorganization,
> > +better media failure reporting, and shrinking.
> > +Second, the different ondisk storage format of the reverse mapping btree
> > +defeats device-level deduplication, because the filesystem requires real
> > +redundancy.
> 
> "defeats device-level deduplication" took me a bit of thought to
> follow what you were trying to describe here.
> 
> I think what you mean is this:
> 
> "Second, the reverse mapping metadata uses a different on-disk
> format for storage. This prevents storage device level deduplication
> from removing redundant metadata the filesystem stores for recovery
> purposes. Hence we will always have two physical copies of the
> metadata we need for recovery in the storage device, regardless of
> the technologies it implements."

Correct.  I'll work that in.

> > +A criticism of adding the secondary index is that it does nothing to improve
> > +the robustness of user data storage itself.
> > +This is a valid point, but adding a new index for file data block checksums
> > +increases write amplification and turns data overwrites into copy-writes, which
> > +age the filesystem prematurely.
> 
> And can come with a significant IO performance penalty.

Right, I shouldn't be assuming that people can draw the connection
between write amplification/aging fs metadata and slow performance.

> > +Checking Extended Attributes
> > +````````````````````````````
> > +
> > +Extended attributes implement a key-value store that enable fragments of data
> > +to be attached to any file.
> > +Both the kernel and userspace can access the keys and values, subject to
> > +namespace and privilege restrictions.
> > +Most typically these fragments are metadata about the file -- origins, security
> > +contexts, user-supplied labels, indexing information, etc.
> > +
> > +Names can be as long as 255 bytes and can exist in several different
> > +namespaces.
> > +Values can be as large as 64KB.
> > +A file's extended attributes are stored in blocks mapped by the attr fork.
> > +The mappings point to leaf blocks, remote value blocks, or dabtree blocks.
> > +Block 0 in the attribute fork is always the top of the structure, but otherwise
> > +each of the three types of blocks can be found at any offset in the attr fork.
> > +Leaf blocks contain attribute key records that point to the name and the value.
> > +Names are always stored elsewhere in the same leaf block.
> > +Values that are less than 3/4 the size of a filesystem block are also stored
> > +elsewhere in the same leaf block.
> > +Remote value blocks contain values that are too large to fit inside a leaf.
> > +If the leaf information exceeds a single filesystem block, a dabtree (also
> > +rooted at block 0) is created to map hashes of the attribute names to leaf
> > +blocks in the attr fork.
> > +
> > +Checking an extended attribute structure is not so straightfoward due to the
> > +lack of separation between attr blocks and index blocks.
> > +Scrub must read each block mapped by the attr fork and ignore the non-leaf
> > +blocks:
> > +
> > +1. Walk the dabtree in the attr fork (if present) to ensure that there are no
> > +   irregularities in the blocks or dabtree mappings that do not point to
> > +   attr leaf blocks.
> > +
> > +2. Walk the blocks of the attr fork looking for leaf blocks.
> > +   For each entry inside a leaf:
> > +
> > +   a. Validate that the name does not contain invalid characters.
> > +
> > +   b. Read the attr value.
> > +      This performs a named lookup of the attr name to ensure the correctness
> > +      of the dabtree.
> > +      If the value is stored in a remote block, this also validates the
> > +      integrity of the remote value block.
> 
> I'm assuming that checking remote xattr blocks involves checking the headers
> are all valid, the overall length is correct, they have valid rmap
> entries, etc?

Yep.  The headers are checked by reading the attr value; and the space
mappings are checked by the attr fork scanner.

> > +Checking and Cross-Referencing Directories
> > +``````````````````````````````````````````
> > +
> > +The filesystem directory tree is a directed acylic graph structure, with files
> > +constituting the nodes, and directory entries (dirents) constituting the edges.
> 
> "with hashed file names consituting the nodes" ?



> > +Directories are a special type of file containing a set of mappings from a
> > +255-byte sequence (name) to an inumber.
> > +These are called directory entries, or dirents for short.
> > +Each directory file must have exactly one directory pointing to the file.
> > +A root directory points to itself.
> > +Directory entries point to files of any type.
> 
> "point to inodes"
> 
> > +Each non-directory file may have multiple directories point to it.
> 
> "Each non directory inode may have multiple dirents point to it"

Both fixed.

> > +
> > +In XFS, directories are implemented as a file containing up to three 32GB
> > +partitions.
> 
> "implemented as offset-based file data"

Fixed.

> > +The first partition contains directory entry data blocks.
> > +Each data block contains variable-sized records associating a user-provided
> > +name with an inumber and, optionally, a file type.
> > +If the directory entry data grows beyond one block, the second partition (which
> > +exists as post-EOF extents) is populated with a block containing free space
> > +information and an index that maps hashes of the dirent names to directory data
> > +blocks in the first partition.
> > +This makes directory name lookups very fast.
> 
> "This is known as a "LEAF 1" format directory."

Good addition!

> > +If this second partition grows beyond one block, the third partition is
> > +populated with a linear array of free space information for faster
> > +expansions.
> 
> s/for faster expansions/to speed up first fit searches when inserting
> new dirents/.
> 
> "This is known as a "NODE" format directory."

Ok.

> > +If the free space has been separated and the second partition grows again
> > +beyond one block, then a dabtree is used to map hashes of dirent names to
> > +directory data blocks.
> 
> The second partition is changed to a dabtree format during NODE
> format conversion - it's just the single root block form. It grows
> as a btree from there by normal split/grow btree operations at that
> point.  Hence I'm not sure this needs to be mentioned as a separate
> step - the node form conversion should mention the second partition
> transitions to the dabtree index format, and it's all obvious from
> there...

I put that nugget in because *I* keep forgetting that there are
directories with a single block in the second partition that doesn't
have the DA3_NODE magicnum. ;)

How about:

"This is known as a 'NODE' format directory.
The second partition contains a dabtree that ranges in size from a
single block to many levels."

> > +Checking a directory is pretty straightfoward:
> > +
> > +1. Walk the dabtree in the second partition (if present) to ensure that there
> > +   are no irregularities in the blocks or dabtree mappings that do not point to
> > +   dirent blocks.
> > +
> > +2. Walk the blocks of the first partition looking for directory entries.
> > +   Each dirent is checked as follows:
> > +
> > +   a. Does the name contain no invalid characters?
> > +
> > +   b. Does the inumber correspond to an actual, allocated inode?
> > +
> > +   c. Does the child inode have a nonzero link count?
> > +
> > +   d. If a file type is included in the dirent, does it match the type of the
> > +      inode?
> > +
> > +   e. If the child is a subdirectory, does the child's dotdot pointer point
> > +      back to the parent?
> > +
> > +   f. If the directory has a second partition, perform a named lookup of the
> > +      dirent name to ensure the correctness of the dabtree.
> > +
> > +3. Walk the free space list in the third partition (if present) to ensure that
> > +   the free spaces it describes are really unused.
> 
> Actaully, they should reflect the largest free space in the given
> block, so this has to be checked against the bests array in the
> given block.
> 
> Also, if we are walking the directory blocks, we should be checking
> the empty ranges for valid tags and sizes, and that the largest 3
> holes in the block are correctly ordered in the bests array, too.

<nod> I think the code actually does that, but I got lazy with writing
here. :/

> > +
> > +Checking operations involving :ref:`parents <dirparent>` and
> > +:ref:`file link counts <nlinks>` are discussed in more detail in later
> > +sections.
> > +
> > +Checking Directory/Attribute Btrees
> > +^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
> > +
> > +As stated in previous sections, the directory/attribute btree (dabtree) index
> > +maps user-provided names to improve lookup times by avoiding linear scans.
> > +Internally, it maps a 32-bit hash of the name to a block offset within the
> > +appropriate file fork.
> > +
> > +The internal structure of a dabtree closely resembles the btrees that record
> > +fixed-size metadata records -- each dabtree block contains a magic number, a
> > +checksum, sibling pointers, a UUID, a tree level, and a log sequence number.
> > +The format of leaf and node records are the same -- each entry points to the
> > +next level down in the hierarchy, with dabtree node records pointing to dabtree
> > +leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere
> > +in the fork.
> > +
> > +Checking and cross-referencing the dabtree is very similar to what is done for
> > +space btrees:
> > +
> > +- Does the type of data stored in the block match what scrub is expecting?
> > +
> > +- Does the block belong to the owning structure that asked for the read?
> > +
> > +- Do the records fit within the block?
> > +
> > +- Are the records contained inside the block free of obvious corruptions?
> > +
> > +- Are the name hashes in the correct order?
> > +
> > +- Do node pointers within the dabtree point to valid fork offsets for dabtree
> > +  blocks?
> > +
> > +- Do leaf pointers within the dabtree point to valid fork offsets for directory
> > +  or attr leaf blocks?
> > +
> > +- Do child pointers point towards the leaves?
> > +
> > +- Do sibling pointers point across the same level?
> > +
> > +- For each dabtree node record, does the record key accurate reflect the
> > +  contents of the child dabtree block?
> > +
> > +- For each dabtree leaf record, does the record key accurate reflect the
> > +  contents of the directory or attr block?
> 
> This last one is checking that they point to a valid dirent and
> check that the hash of the dirent name actually matches the hash
> that points to it?

Yes.

--D

> Cheers,
> 
> Dave.
> -- 
> Dave Chinner
> david@fromorbit.com

[PATCH 06/14] xfs: document how online fsck deals with eventual consistency

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Writes to an XFS filesystem employ an eventual consistency update model
to break up complex multistep metadata updates into small chained
transactions.  This is generally good for performance and scalability
because XFS doesn't need to prepare for enormous transactions, but it
also means that online fsck must be careful not to attempt a fsck action
unless it can be shown that there are no other threads processing a
transaction chain.  This part of the design documentation covers the
thinking behind the consistency model and how scrub deals with it.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  307 ++++++++++++++++++++
 1 file changed, 307 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index f45bf97fa9c4..7b783a0e85b9 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -1443,3 +1443,310 @@ This step is critical for enabling system administrator to monitor the status
 of the filesystem and the progress of any repairs.
 For developers, it is a useful means to judge the efficacy of error detection
 and correction in the online and offline checking tools.
+
+Eventual Consistency vs. Online Fsck
+------------------------------------
+
+Midway through the development of online scrubbing, the fsstress tests
+uncovered a misinteraction between online fsck and compound transaction chains
+created by other writer threads that resulted in false reports of metadata
+inconsistency.
+The root cause of these reports is the eventual consistency model introduced by
+the expansion of deferred work items and compound transaction chains when
+reverse mapping and reflink were introduced.
+
+Originally, transaction chains were added to XFS to avoid deadlocks when
+unmapping space from files.
+Deadlock avoidance rules require that AGs only be locked in increasing order,
+which makes it impossible (say) to use a single transaction to free a space
+extent in AG 7 and then try to free a now superfluous block mapping btree block
+in AG 3.
+To avoid these kinds of deadlocks, XFS creates Extent Freeing Intent (EFI) log
+items to commit to freeing some space in one transaction while deferring the
+actual metadata updates to a fresh transaction.
+The transaction sequence looks like this:
+
+1. The first transaction contains a physical update to the file's block mapping
+   structures to remove the mapping from the btree blocks.
+   It then attaches to the in-memory transaction an action item to schedule
+   deferred freeing of space.
+   Concretely, each transaction maintains a list of ``struct
+   xfs_defer_pending`` objects, each of which maintains a list of ``struct
+   xfs_extent_free_item`` objects.
+   Returning to the example above, the action item tracks the freeing of both
+   the unmapped space from AG 7 and the block mapping btree (BMBT) block from
+   AG 3.
+   Deferred frees recorded in this manner are committed in the log by creating
+   an EFI log item from the ``struct xfs_extent_free_item`` object and
+   attaching the log item to the transaction.
+   When the log is persisted to disk, the EFI item is written into the ondisk
+   transaction record.
+   EFIs can list up to 16 extents to free, all sorted in AG order.
+
+2. The second transaction contains a physical update to the free space btrees
+   of AG 3 to release the former BMBT block and a second physical update to the
+   free space btrees of AG 7 to release the unmapped file space.
+   Observe that the the physical updates are resequenced in the correct order
+   when possible.
+   Attached to the transaction is a an extent free done (EFD) log item.
+   The EFD contains a pointer to the EFI logged in transaction #1 so that log
+   recovery can tell if the EFI needs to be replayed.
+
+If the system goes down after transaction #1 is written back to the filesystem
+but before #2 is committed, a scan of the filesystem metadata would show
+inconsistent filesystem metadata because there would not appear to be any owner
+of the unmapped space.
+Happily, log recovery corrects this inconsistency for us -- when recovery finds
+an intent log item but does not find a corresponding intent done item, it will
+reconstruct the incore state of the intent item and finish it.
+In the example above, the log must replay both frees described in the recovered
+EFI to complete the recovery phase.
+
+There are two subtleties to XFS' transaction chaining strategy to consider.
+The first is that log items must be added to a transaction in the correct order
+to prevent conflicts with principal objects that are not held by the
+transaction.
+In other words, all per-AG metadata updates for an unmapped block must be
+completed before the last update to free the extent, and extents should not
+be reallocated until that last update commits to the log.
+The second subtlety comes from the fact that AG header buffers are (usually)
+released between each transaction in a chain.
+This means that other threads can observe an AG in an intermediate state,
+but as long as the first subtlety is handled, this should not affect the
+correctness of filesystem operations.
+Unmounting the filesystem flushes all pending work to disk, which means that
+offline fsck never sees the temporary inconsistencies caused by deferred work
+item processing.
+In this manner, XFS employs a form of eventual consistency to avoid deadlocks
+and increase parallelism.
+
+During the design phase of the reverse mapping and reflink features, it was
+decided that it was impractical to cram all the reverse mapping updates for a
+single filesystem change into a single transaction because a single file
+mapping operation can explode into many small updates:
+
+* The block mapping update itself
+* A reverse mapping update for the block mapping update
+* Fixing the freelist
+* A reverse mapping update for the freelist fix
+
+* A shape change to the block mapping btree
+* A reverse mapping update for the btree update
+* Fixing the freelist (again)
+* A reverse mapping update for the freelist fix
+
+* An update to the reference counting information
+* A reverse mapping update for the refcount update
+* Fixing the freelist (a third time)
+* A reverse mapping update for the freelist fix
+
+* Freeing any space that was unmapped and not owned by any other file
+* Fixing the freelist (a fourth time)
+* A reverse mapping update for the freelist fix
+
+* Freeing the space used by the block mapping btree
+* Fixing the freelist (a fifth time)
+* A reverse mapping update for the freelist fix
+
+Free list fixups are not usually needed more than once per AG per transaction
+chain, but it is theoretically possible if space is very tight.
+For copy-on-write updates this is even worse, because this must be done once to
+remove the space from a staging area and again to map it into the file!
+
+To deal with this explosion in a calm manner, XFS expands its use of deferred
+work items to cover most reverse mapping updates and all refcount updates.
+This reduces the worst case size of transaction reservations by breaking the
+work into a long chain of small updates, which increases the degree of eventual
+consistency in the system.
+Again, this generally isn't a problem because XFS orders its deferred work
+items carefully to avoid resource reuse conflicts between unsuspecting threads.
+
+However, online fsck changes the rules -- remember that although physical
+updates to per-AG structures are coordinated by locking the buffers for AG
+headers, buffer locks are dropped between transactions.
+Once scrub acquires resources and takes locks for a data structure, it must do
+all the validation work without releasing the lock.
+If the main lock for a space btree is an AG header buffer lock, scrub may have
+interrupted another thread that is midway through finishing a chain.
+For example, if a thread performing a copy-on-write has completed a reverse
+mapping update but not the corresponding refcount update, the two AG btrees
+will appear inconsistent to scrub and an observation of corruption will be
+recorded.  This observation will not be correct.
+If a repair is attempted in this state, the results will be catastrophic!
+
+Several solutions to this problem were evaluated upon discovery of this flaw:
+
+1. Add a higher level lock to allocation groups and require writer threads to
+   acquire the higher level lock in AG order before making any changes.
+   This would be very difficult to implement in practice because it is
+   difficult to determine which locks need to be obtained, and in what order,
+   without simulating the entire operation.
+   Performing a dry run of a file operation to discover necessary locks would
+   make the filesystem very slow.
+
+2. Make the deferred work coordinator code aware of consecutive intent items
+   targeting the same AG and have it hold the AG header buffers locked across
+   the transaction roll between updates.
+   This would introduce a lot of complexity into the coordinator since it is
+   only loosely coupled with the actual deferred work items.
+   It would also fail to solve the problem because deferred work items can
+   generate new deferred subtasks, but all subtasks must be complete before
+   work can start on a new sibling task.
+
+3. Teach online fsck to walk all transactions waiting for whichever lock(s)
+   protect the data structure being scrubbed to look for pending operations.
+   The checking and repair operations must factor these pending operations into
+   the evaluations being performed.
+   This solution is a nonstarter because it is *extremely* invasive to the main
+   filesystem.
+
+4. Recognize that only online fsck has this requirement of total consistency
+   of AG metadata, and that online fsck should be relatively rare as compared
+   to filesystem change operations.
+   For each AG, maintain a sloppy count of intent items targetting that AG.
+   When online fsck wants to examine an AG, it should lock the AG header
+   buffers to quiesce all transaction chains that want to modify that AG, and
+   only proceed with the scrub if the count is zero.
+   In other words, scrub only proceeds if it can lock the AG header buffers and
+   there can't possibly be any intents in progress.
+   This may lead to fairness and starvation issues, but regular filesystem
+   updates take precedence over online fsck activity.
+
+Intent Drains
+`````````````
+
+The fourth solution is implemented in the current iteration of online fsck,
+with percpu counters providing the "sloppy" counter.
+Updates to the percpu counter from normal writer threads are very fast, which
+is good for maintaining runtime performance.
+
+There are two key properties to the drain mechanism.
+First, the counter is incremented when a deferred work item is *queued* to a
+transaction, and it is decremented after the associated intent done log item is
+*committed* to another transaction.
+The second property is that deferred work can be added to a transaction without
+holding an AG header lock, but per-AG work items cannot be marked done without
+locking that AG header buffer to log the physical updates and the intent done
+log item.
+The first property enables scrub to yield to running transaction chains, which
+is an explicit deprioritization of online fsck to benefit file operations.
+The second property of the drain is key to the correct coordination of scrub,
+since scrub will always be able to decide if a conflict is possible.
+
+For regular filesystem code, the drain works as follows:
+
+1. Add a deferred item to a transaction.
+
+2. The deferred item manager calls the ``->add_item`` method of the item.
+
+3. The ``->add_item`` implementation calls ``xfs_drain_bump`` to increase the
+   sloppy counter.
+
+4. When the deferred item manager wants to finish the defeferred work, it calls
+   ``->finish_item`` to complete it.
+
+5. The ``->finish_item`` implementation logs some changes and calls
+   ``xfs_drain_drop`` to decrease the sloppy counter and wake up any threads
+   waiting on the drain.
+
+6. The subtransaction commits, which unlocks the resource associated with the
+   intent item.
+
+For scrub, the drain works as follows:
+
+1. Lock the resource(s) associated with the metadata being scrubbed.
+   For example, a scan of the refcount btree would lock the AGI and AGF header
+   buffers.
+
+2. If the sloppy counter is zero (``xfs_drain_busy`` returns false), there are
+   no chains in progress and the operation may proceed.
+
+3. Otherwise, release the resources grabbed in step 1.
+
+4. Wait for the intent counter to reach zero (``xfs_drain_intents``), then go
+   back to step 1 unless a signal has been caught.
+
+To avoid polling in step 4, the drain provides a waitqueue for scrub threads to
+be woken up whenever the intent count drops.
+
+The proposed patchset is the
+`scrub intent drain series
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-drain-intents>`_.
+
+.. _jump_labels:
+
+Static Keys (aka Jump Label Patching)
+`````````````````````````````````````
+
+Online fsck for XFS separates the regular filesystem from the checking and
+repair code as much as possible.
+However, there are a few parts of online fsck (such as the intent drains, and
+later, live update hooks) where it is useful for the online fsck code to know
+what's going on in the rest of the filesystem.
+Since it is not expected that online fsck will be constantly running in the
+background, it is very important to minimize the runtime overhead imposed by
+these hooks when online fsck is compiled into the kernel but not actively
+running on behalf of userspace.
+Taking locks in the hot path of a writer thread to access a data structure only
+to find that no further action is necessary is expensive -- on the author's
+computer, this have an overhead of 40-50ns per access.
+Fortunately, the kernel supports dynamic code patching, which enables XFS to
+replace a static branch to hook code with ``nop`` sleds when online fsck isn't
+running.
+This sled has an overhead of however long it takes the instruction decoder to
+skip past the sled, which seems to be on the order of less than 1ns and
+does not access memory outside of instruction fetching.
+
+When online fsck enables the static key, the sled is replaced with an
+unconditional branch to call the hook code.
+The switchover is quite expensive (~22000ns) but is paid entirely by the
+program that invoked online fsck, and can be amortized if multiple threads
+enter online fsck at the same time, or if multiple filesystems are being
+checked at the same time.
+Changing the branch direction requires taking the CPU hotplug lock, and since
+CPU initialization requires memory allocation, online fsck must be careful not
+to change a static key while holding any locks or resources that could be
+accessed in the memory reclaim paths.
+To minimize contention on the CPU hotplug lock, care should be taken not to
+enable or disable static keys unnecessarily.
+
+Because static keys are intended to minimize hook overhead for regular
+filesystem operations when xfs_scrub is not running, the intended usage
+patterns are as follows:
+
+- The hooked part of XFS should declare a static-scoped static key that
+  defaults to false.
+  The ``DEFINE_STATIC_KEY_FALSE`` macro takes care of this.
+  The static key itself should be declared as a ``static`` variable.
+
+- When deciding to invoke code that's only used by scrub, the regular
+  filesystem should call the ``static_branch_unlikely`` predicate to avoid the
+  scrub-only hook code if the static key is not enabled.
+
+- The regular filesystem should export helper functions that call
+  ``static_branch_inc`` to enable and ``static_branch_dec`` to disable the
+  static key.
+  Wrapper functions make it easy to compile out the relevant code if the kernel
+  distributor turns off online fsck at build time.
+
+- Scrub functions wanting to turn on scrub-only XFS functionality should call
+  the ``xchk_fshooks_enable`` from the setup function to enable a specific
+  hook.
+  This must be done before obtaining any resources that are used by memory
+  reclaim.
+  Callers had better be sure they really need the functionality gated by the
+  static key; the ``TRY_HARDER`` flag is useful here.
+
+Online scrub has resource acquisition helpers (e.g. ``xchk_perag_lock``) to
+handle locking AGI and AGF buffers for all scrubber functions.
+If it detects a conflict between scrub and the running transactions, it will
+try to wait for intents to complete.
+If the caller of the helper has not enabled the static key, the helper will
+return -EDEADLOCK, which should result in the scrub being restarted with the
+``TRY_HARDER`` flag set.
+The scrub setup function should detect that flag, enable the static key, and
+try the scrub again.
+Scrub teardown disables all static keys obtained by ``xchk_fshooks_enable``.
+
+For more information, please see the kernel documentation of
+Documentation/staging/static-keys.rst.

[PATCH 07/14] xfs: document pageable kernel memory

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Add a discussion of pageable kernel memory, since online fsck needs
quite a bit more memory than most other parts of the filesystem to stage
records and other information.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  490 ++++++++++++++++++++
 1 file changed, 490 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 7b783a0e85b9..2957aefb983a 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -383,6 +383,8 @@ Algorithms") of Srinivasan.
 However, any data structure builder that maintains a resource lock for the
 duration of the repair is *always* an offline algorithm.
 
+.. _secondary_metadata:
+
 Secondary Metadata
 ``````````````````
 
@@ -1750,3 +1752,491 @@ Scrub teardown disables all static keys obtained by ``xchk_fshooks_enable``.
 
 For more information, please see the kernel documentation of
 Documentation/staging/static-keys.rst.
+
+.. _xfile:
+
+Pageable Kernel Memory
+----------------------
+
+Demonstrations of the first few prototypes of online repair revealed new
+technical requirements that were not originally identified.
+For the first demonstration, the code walked whatever filesystem
+metadata it needed to synthesize new records and inserted records into a new
+btree as it found them.
+This was subpar since any additional corruption or runtime errors encountered
+during the walk would shut down the filesystem.
+After remount, the blocks containing the half-rebuilt data structure would not
+be accessible until another repair was attempted.
+Solving the problem of half-rebuilt data structures will be discussed in the
+next section.
+
+For the second demonstration, the synthesized records were instead stored in
+kernel slab memory.
+Doing so enabled online repair to abort without writing to the filesystem if
+the metadata walk failed, which prevented online fsck from making things worse.
+However, even this approach needed improving upon.
+
+There are four reasons why traditional Linux kernel memory management isn't
+suitable for storing large datasets:
+
+1. Although it is tempting to allocate a contiguous block of memory to create a
+   C array, this cannot easily be done in the kernel because it cannot be
+   relied upon to allocate multiple contiguous memory pages.
+
+2. While disparate physical pages can be virtually mapped together, installed
+   memory might still not be large enough to stage the entire record set in
+   memory while constructing a new btree.
+
+3. To overcome these two difficulties, the implementation was adjusted to use
+   doubly linked lists, which means every record object needed two 64-bit list
+   head pointers, which is a lot of overhead.
+
+4. Kernel memory is pinned, which can drive the system out of memory, leading
+   to OOM kills of unrelated processes.
+
+For the third iteration, attention swung back to the possibility of using
+byte-indexed array-like storage to reduce the overhead of in-memory records.
+At any given time, online repair does not need to keep the entire record set in
+memory, which means that individual records can be paged out.
+Creating new temporary files in the XFS filesystem to store intermediate data
+was explored and rejected for some types of repairs because a filesystem with
+compromised space and inode metadata should never be used to fix compromised
+space or inode metadata.
+However, the kernel already has a facility for byte-addressable and pageable
+storage: shmfs.
+In-kernel graphics drivers (most notably i915) take advantage of shmfs files
+to store intermediate data that doesn't need to be in memory at all times, so
+that usage precedent is already established.
+Hence, the ``xfile`` was born!
+
+xfile Access Models
+```````````````````
+
+A survey of the intended uses of xfiles suggested these use cases:
+
+1. Arrays of fixed-sized records (space management btrees, directory and
+   extended attribute entries)
+
+2. Sparse arrays of fixed-sized records (quotas and link counts)
+
+3. Large binary objects (BLOBs) of variable sizes (directory and extended
+   attribute names and values)
+
+4. Staging btrees in memory (reverse mapping btrees)
+
+5. Arbitrary contents (realtime space management)
+
+To support the first four use cases, high level data structures wrap the xfile
+to share functionality between online fsck functions.
+The rest of this section discusses the interfaces that the xfile presents to
+four of those five higher level data structures.
+The fifth use case is discussed in the :ref:`realtime summary <rtsummary>` case
+study.
+
+The most general storage interface supported by the xfile enables the reading
+and writing of arbitrary quantities of data at arbitrary offsets in the xfile.
+This capability is provided by ``xfile_pread`` and ``xfile_pwrite`` functions,
+which behave similarly to their userspace counterparts.
+XFS is very record-based, which suggests that the ability to load and store
+complete records is important.
+To support these cases, a pair of ``xfile_obj_load`` and ``xfile_obj_store``
+functions are provided to read and persist objects into an xfile.
+They are internally the same as pread and pwrite, except that they treat any
+error as an out of memory error.
+For online repair, squashing error conditions in this manner is an acceptable
+behavior because the only reaction is to abort the operation back to userspace.
+All five xfile usecases can be serviced by these four functions.
+
+However, no discussion of file access idioms is complete without answering the
+question, "But what about mmap?"
+It would be *much* more convenient if kernel code could access pageable kernel
+memory with pointers, just like userspace code does with regular memory.
+Like any other filesystem that uses the page cache, reads and writes of xfile
+data lock the cache page and map it into the kernel address space for the
+duration of the operation.
+Unfortunately, shmfs can only write a file page to the swap device if the page
+is unmapped and unlocked, which means the xfile risks causing OOM problems
+unless it is careful not to pin too many pages.
+Therefore, the xfile steers most of its users towards programmatic access so
+that backing pages are not kept locked in memory for longer than is necessary.
+However, for callers performing quick linear scans of xfile data,
+``xfile_obj_get_page`` and ``xfile_obj_put_page`` functions are provided to pin
+a page in memory.
+So far, the only code to use these functions are the xfarray :ref:`sorting
+<xfarray_sort>` algorithms.
+
+xfile Access Coordination
+`````````````````````````
+
+For security reasons, xfiles must be owned privately by the kernel.
+They are marked ``S_PRIVATE`` to prevent interference from the security system,
+must never be mapped into process file descriptor tables, and their pages must
+never be mapped into userspace processes.
+
+To avoid locking recursion issues with the VFS, all accesses to the shmfs file
+are performed by manipulating the page cache directly.
+xfile writes call the ``->write_begin`` and ``->write_end`` functions of the
+xfile's address space to grab writable pages, copy the caller's buffer into the
+page, and release the pages.
+xfile reads call ``shmem_read_mapping_page_gfp`` to grab pages directly before
+copying the contents into the caller's buffer.
+In other words, xfiles ignore the VFS read and write code paths to avoid
+having to create a dummy ``struct kiocb`` and to avoid taking inode and
+freeze locks.
+
+If an xfile is shared between threads to stage repairs, the caller must provide
+its own locks to coordinate access.
+
+.. _xfarray:
+
+Arrays of Fixed-Sized Records
+`````````````````````````````
+
+In XFS, each type of indexed space metadata (free space, inodes, reference
+counts, file fork space, and reverse mappings) consists of a set of fixed-size
+records indexed with a classic B+ tree.
+Directories have a set of fixed-size dirent records that point to the names,
+and extended attributes have a set of fixed-size attribute keys that point to
+names and values.
+Quota counters and file link counters index records with numbers.
+During a repair, scrub needs to stage new records during the gathering step and
+retrieve them during the btree building step.
+
+Although this requirement can be satisfied by calling the read and write
+methods of the xfile directly, it is simpler for callers for there to be a
+higher level abstraction to take care of computing array offsets, to provide
+iterator functions, and to deal with sparse records and sorting.
+The ``xfarray`` abstraction presents a linear array for fixed-size records atop
+the byte-accessible xfile.
+
+.. _xfarray_access_patterns:
+
+Array Access Patterns
+^^^^^^^^^^^^^^^^^^^^^
+
+Array access patterns in online fsck tend to fall into three categories.
+Iteration of records is assumed to be necessary for all cases and will be
+covered in the next section.
+
+The first type of caller handles records that are indexed by position.
+Gaps may exist between records, and a record may be updated multiple times
+during the collection step.
+In other words, these callers want a sparse linearly addressed table file.
+The typical use case are quota records or file link count records.
+Access to array elements is performed programmatically via ``xfarray_load`` and
+``xfarray_store`` functions, which wrap the similarly-named xfile functions to
+provide loading and storing of array elements at arbitrary array indices.
+Gaps are defined to be null records, and null records are defined to be a
+sequence of all zero bytes.
+Null records are detected by calling ``xfarray_element_is_null``.
+They are created either by calling ``xfarray_unset`` to null out an existing
+record or by never storing anything to an array index.
+
+The second type of caller handles records that are not indexed by position
+and do not require multiple updates to a record.
+The typical use case here is rebuilding space btrees and key/value btrees.
+These callers can add records to the array without caring about array indices
+via the ``xfarray_append`` function, which stores a record at the end of the
+array.
+For callers that require records to be presentable in a specific order (e.g.
+rebuilding btree data), the ``xfarray_sort`` function can arrange the sorted
+records; this function will be covered later.
+
+The third type of caller is a bag, which is useful for counting records.
+The typical use case here is constructing space extent reference counts from
+reverse mapping information.
+Records can be put in the bag in any order, they can be removed from the bag
+at any time, and uniqueness of records is left to callers.
+The ``xfarray_store_anywhere`` function is used to insert a record in any
+null record slot in the bag; and the ``xfarray_unset`` function removes a
+record from the bag.
+
+The proposed patchset is the
+`big in-memory array
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=big-array>`_.
+
+Iterating Array Elements
+^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most users of the xfarray require the ability to iterate the records stored in
+the array.
+Callers can probe every possible array index with the following:
+
+.. code-block:: c
+
+	xfarray_idx_t i;
+	foreach_xfarray_idx(array, i) {
+	    xfarray_load(array, i, &rec);
+
+	    /* do something with rec */
+	}
+
+All users of this idiom must be prepared to handle null records or must already
+know that there aren't any.
+
+For xfarray users that want to iterate a sparse array, the ``xfarray_iter``
+function ignores indices in the xfarray that have never been written to by
+calling ``xfile_seek_data`` (which internally uses ``SEEK_DATA``) to skip areas
+of the array that are not populated with memory pages.
+Once it finds a page, it will skip the zeroed areas of the page.
+
+.. code-block:: c
+
+	xfarray_idx_t i = XFARRAY_CURSOR_INIT;
+	while ((ret = xfarray_iter(array, &i, &rec)) == 1) {
+	    /* do something with rec */
+	}
+
+.. _xfarray_sort:
+
+Sorting Array Elements
+^^^^^^^^^^^^^^^^^^^^^^
+
+During the fourth demonstration of online repair, a community reviewer remarked
+that for performance reasons, online repair ought to load batches of records
+into btree record blocks instead of inserting records into a new btree one at a
+time.
+The btree insertion code in XFS is responsible for maintaining correct ordering
+of the records, so naturally the xfarray must also support sorting the record
+set prior to bulk loading.
+
+The sorting algorithm used in the xfarray is actually a combination of adaptive
+quicksort and a heapsort subalgorithm in the spirit of
+`Sedgewick <https://algs4.cs.princeton.edu/23quicksort/>`_ and
+`pdqsort <https://github.com/orlp/pdqsort>`_, with customizations for the Linux
+kernel.
+To sort records in a reasonably short amount of time, ``xfarray`` takes
+advantage of the binary subpartitioning offered by quicksort, but it also uses
+heapsort to hedge aginst performance collapse if the chosen quicksort pivots
+are poor.
+Both algorithms are (in general) O(n * lg(n)), but there is a wide performance
+gulf between the two implementations.
+
+The Linux kernel already contains a reasonably fast implementation of heapsort.
+It only operates on regular C arrays, which limits the scope of its usefulness.
+There are two key places where the xfarray uses it:
+
+* Sorting any record subset backed by a single xfile page.
+
+* Loading a small number of xfarray records from potentially disparate parts
+  of the xfarray into a memory buffer, and sorting the buffer.
+
+In other words, ``xfarray`` uses heapsort to constrain the nested recursion of
+quicksort, thereby mitigating quicksort's worst runtime behavior.
+
+Choosing a quicksort pivot is a tricky business.
+A good pivot splits the set to sort in half, leading to the divide and conquer
+behavior that is crucial to  O(n * lg(n)) performance.
+A poor pivot barely splits the subset at all, leading to O(n\ :sup:`2`)
+runtime.
+The xfarray sort routine tries to avoid picking a bad pivot by sampling nine
+records into a memory buffer and using the kernel heapsort to identify the
+median of the nine.
+
+Most modern quicksort implementations employ Tukey's "ninther" to select a
+pivot from a classic C array.
+Typical ninther implementations pick three unique triads of records, sort each
+of the triads, and then sort the middle value of each triad to determine the
+ninther value.
+As stated previously, however, xfile accesses are not entirely cheap.
+It turned out to be much more performant to read the nine elements into a
+memory buffer, run the kernel's in-memory heapsort on the buffer, and choose
+the 4th element of that buffer as the pivot.
+Tukey's ninthers are described in J. W. Tukey, `The ninther, a technique for
+low-effort robust (resistant) location in large samples`, in *Contributions to
+Survey Sampling and Applied Statistics*, edited by H. David, (Academic Press,
+1978), pp. 251–257.
+
+The partitioning of quicksort is fairly textbook -- rearrange the record
+subset around the pivot, then set up the current and next stack frames to
+sort with the larger and the smaller halves of the pivot, respectively.
+This keeps the stack space requirements to log2(record count).
+
+As a final performance optimization, the hi and lo scanning phase of quicksort
+keeps examined xfile pages mapped in the kernel for as long as possible to
+reduce map/unmap cycles.
+Surprisingly, this reduces overall sort runtime by nearly half again after
+accounting for the application of heapsort directly onto xfile pages.
+
+Blob Storage
+````````````
+
+Extended attributes and directories add an additional requirement for staging
+records: arbitrary byte sequences of finite length.
+Each directory entry record needs to store entry name,
+and each extended attribute needs to store both the attribute name and value.
+The names, keys, and values can consume a large amount of memory, so the
+``xfblob`` abstraction was created to simplify management of these blobs
+atop an xfile.
+
+Blob arrays provide ``xfblob_load`` and ``xfblob_store`` functions to retrieve
+and persist objects.
+The store function returns a magic cookie for every object that it persists.
+Later, callers provide this cookie to the ``xblob_load`` to recall the object.
+The ``xfblob_free`` function frees a specific blob, and the ``xfblob_truncate``
+function frees them all because compaction is not needed.
+
+The details of repairing directories and extended attributes will be discussed
+in a subsequent section about atomic extent swapping.
+However, it should be noted that these repair functions only use blob storage
+to cache a small number of entries before adding them to a temporary ondisk
+file, which is why compaction is not required.
+
+The proposed patchset is at the start of the
+`extended attribute repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-xattrs>`_ series.
+
+.. _xfbtree:
+
+In-Memory B+Trees
+`````````````````
+
+The chapter about :ref:`secondary metadata<secondary_metadata>` mentioned that
+checking and repairing of secondary metadata commonly requires coordination
+between a live metadata scan of the filesystem and writer threads that are
+updating that metadata.
+Keeping the scan data up to date requires requires the ability to propagate
+metadata updates from the filesystem into the data being collected by the scan.
+This *can* be done by appending concurrent updates into a separate log file and
+applying them before writing the new metadata to disk, but this leads to
+unbounded memory consumption if the rest of the system is very busy.
+Another option is to skip the side-log and commit live updates from the
+filesystem directly into the scan data, which trades more overhead for a lower
+maximum memory requirement.
+In both cases, the data structure holding the scan results must support indexed
+access to perform well.
+
+Given that indexed lookups of scan data is required for both strategies, online
+fsck employs the second strategy of committing live updates directly into
+scan data.
+Because xfarrays are not indexed and do not enforce record ordering, they
+are not suitable for this task.
+Conveniently, however, XFS has a library to create and maintain ordered reverse
+mapping records: the existing rmap btree code!
+If only there was a means to create one in memory.
+
+Recall that the :ref:`xfile <xfile>` abstraction represents memory pages as a
+regular file, which means that the kernel can create byte or block addressable
+virtual address spaces at will.
+The XFS buffer cache specializes in abstracting IO to block-oriented  address
+spaces, which means that adaptation of the buffer cache to interface with
+xfiles enables reuse of the entire btree library.
+Btrees built atop an xfile are collectively known as ``xfbtrees``.
+The next few sections describe how they actually work.
+
+The proposed patchset is the
+`in-memory btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=in-memory-btrees>`_
+series.
+
+Using xfiles as a Buffer Cache Target
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Two modifications are necessary to support xfiles as a buffer cache target.
+The first is to make it possible for the ``struct xfs_buftarg`` structure to
+host the ``struct xfs_buf`` rhashtable, because normally those are held by a
+per-AG structure.
+The second change is to modify the buffer ``ioapply`` function to "read" cached
+pages from the xfile and "write" cached pages back to the xfile.
+Multiple access to individual buffers is controlled by the ``xfs_buf`` lock,
+since the xfile does not provide any locking on its own.
+With this adaptation in place, users of the xfile-backed buffer cache use
+exactly the same APIs as users of the disk-backed buffer cache.
+The separation between xfile and buffer cache implies higher memory usage since
+they do not share pages, but this property could some day enable transactional
+updates to an in-memory btree.
+Today, however, it simply eliminates the need for new code.
+
+Space Management with an xfbtree
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Space management for an xfile is very simple -- each btree block is one memory
+page in size.
+These blocks use the same header format as an on-disk btree, but the in-memory
+block verifiers ignore the checksums, assuming that xfile memory is no more
+corruption-prone than regular DRAM.
+Reusing existing code here is more important than absolute memory efficiency.
+
+The very first block of an xfile backing an xfbtree contains a header block.
+The header describes the owner, height, and the block number of the root
+xfbtree block.
+
+To allocate a btree block, use ``xfile_seek_data`` to find a gap in the file.
+If there are no gaps, create one by extending the length of the xfile.
+Preallocate space for the block with ``xfile_prealloc``, and hand back the
+location.
+To free an xfbtree block, use ``xfile_discard`` (which internally uses
+``FALLOC_FL_PUNCH_HOLE``) to remove the memory page from the xfile.
+
+Populating an xfbtree
+^^^^^^^^^^^^^^^^^^^^^
+
+An online fsck function that wants to create an xfbtree should proceed as
+follows:
+
+1. Call ``xfile_create`` to create an xfile.
+
+2. Call ``xfs_alloc_memory_buftarg`` to create a buffer cache target structure
+   pointing to the xfile.
+
+3. Pass the buffer cache target, buffer ops, and other information to
+   ``xfbtree_create`` to write an initial tree header and root block to the
+   xfile.
+   Each btree type should define a wrapper that passes necessary arguments to
+   the creation function.
+   For example, rmap btrees define ``xfs_rmapbt_mem_create`` to take care of
+   all the necessary details for callers.
+   A ``struct xfbtree`` object will be returned.
+
+4. Pass the xfbtree object to the btree cursor creation function for the
+   btree type.
+   Following the example above, ``xfs_rmapbt_mem_cursor`` takes care of this
+   for callers.
+
+5. Pass the btree cursor to the regular btree functions to make queries against
+   and to update the in-memory btree.
+   For example, a btree cursor for an rmap xfbtree can be passed to the
+   ``xfs_rmap_*`` functions just like any other btree cursor.
+   See the :ref:`next section<xfbtree_commit>` for information on dealing with
+   xfbtree updates that are logged to a transaction.
+
+6. When finished, delete the btree cursor, destroy the xfbtree object, free the
+   buffer target, and the destroy the xfile to release all resources.
+
+.. _xfbtree_commit:
+
+Committing Logged xfbtree Buffers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Although it is a clever hack to reuse the rmap btree code to handle the staging
+structure, the ephemeral nature of the in-memory btree block storage presents
+some challenges of its own.
+The XFS transaction manager must not commit buffer log items for buffers backed
+by an xfile because the log format does not understand updates for devices
+other than the data device.
+An ephemeral xfbtree probably will not exist by the time the AIL checkpoints
+log transactions back into the filesystem, and certainly won't exist during
+log recovery.
+For these reasons, any code updating an xfbtree in transaction context must
+remove the buffer log items from the transaction and write the updates into the
+backing xfile before committing or cancelling the transaction.
+
+The ``xfbtree_trans_commit`` and ``xfbtree_trans_cancel`` functions implement
+this functionality as follows:
+
+1. Find each buffer log item whose buffer targets the xfile.
+
+2. Record the dirty/ordered status of the log item.
+
+3. Detach the log item from the buffer.
+
+4. Queue the buffer to a special delwri list.
+
+5. Clear the transaction dirty flag if the only dirty log items were the ones
+   that were detached in step 3.
+
+6. Submit the delwri list to commit the changes to the xfile, if the updates
+   are being committed.
+
+After removing xfile logged buffers from the transaction in this manner, the
+transaction can be committed or cancelled.

[PATCH 08/14] xfs: document btree bulk loading

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Add a discussion of the btree bulk loading code, which makes it easy to
take an in-memory recordset and write it out to disk in an efficient
manner.  This also enables atomic switchover from the old to the new
structure with minimal potential for leaking the old blocks.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  632 ++++++++++++++++++++
 1 file changed, 632 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 2957aefb983a..c41f089549a0 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -2240,3 +2240,635 @@ this functionality as follows:
 
 After removing xfile logged buffers from the transaction in this manner, the
 transaction can be committed or cancelled.
+
+Bulk Loading of Ondisk B+Trees
+------------------------------
+
+As mentioned previously, early iterations of online repair built new btree
+structures by creating a new btree and adding observations individually.
+Loading a btree one record at a time had a slight advantage of not requiring
+the incore records to be sorted prior to commit, but was very slow and leaked
+blocks if the system went down during a repair.
+Loading records one at a time also meant that repair could not control the
+loading factor of the blocks in the new btree.
+
+Fortunately, the venerable ``xfs_repair`` tool had a more efficient means for
+rebuilding a btree index from a collection of records -- bulk btree loading.
+This was implemented rather inefficiently code-wise, since ``xfs_repair``
+had separate copy-pasted implementations for each btree type.
+
+To prepare for online fsck, each of the four bulk loaders were studied, notes
+were taken, and the four were refactored into a single generic btree bulk
+loading mechanism.
+Those notes in turn have been refreshed and are presented below.
+
+Geometry Computation
+````````````````````
+
+The zeroth step of bulk loading is to assemble the entire record set that will
+be stored in the new btree, and sort the records.
+Next, call ``xfs_btree_bload_compute_geometry`` to compute the shape of the
+btree from the record set, the type of btree, and any load factor preferences.
+This information is required for resource reservation.
+
+First, the geometry computation computes the minimum and maximum records that
+will fit in a leaf block from the size of a btree block and the size of the
+block header.
+Roughly speaking, the maximum number of records is::
+
+        maxrecs = (block_size - header_size) / record_size
+
+The XFS design specifies that btree blocks should be merged when possible,
+which means the minimum number of records is half of maxrecs::
+
+        minrecs = maxrecs / 2
+
+The next variable to determine is the desired loading factor.
+This must be at least minrecs and no more than maxrecs.
+Choosing minrecs is undesirable because it wastes half the block.
+Choosing maxrecs is also undesirable because adding a single record to each
+newly rebuilt leaf block will cause a tree split, which causes a noticeable
+drop in performance immediately afterwards.
+The default loading factor was chosen to be 75% of maxrecs, which provides a
+reasonably compact structure without any immediate split penalties.
+If space is tight, the loading factor will be set to maxrecs to try to avoid
+running out of space::
+
+        leaf_load_factor = enough space ? (maxrecs + minrecs) / 2 : maxrecs
+
+Load factor is computed for btree node blocks using the combined size of the
+btree key and pointer as the record size::
+
+        maxrecs = (block_size - header_size) / (key_size + ptr_size)
+        minrecs = maxrecs / 2
+        node_load_factor = enough space ? (maxrecs + minrecs) / 2 : maxrecs
+
+Once that's done, the number of leaf blocks required to store the record set
+can be computed as::
+
+        leaf_blocks = ceil(record_count / leaf_load_factor)
+
+The number of node blocks needed to point to the next level down in the tree
+is computed as::
+
+        n_blocks = (n == 0 ? leaf_blocks : node_blocks[n])
+        node_blocks[n + 1] = ceil(n_blocks / node_load_factor)
+
+The entire computation is performed recursively until the current level only
+needs one block.
+The resulting geometry is as follows:
+
+- For AG-rooted btrees, this level is the root level, so the height of the new
+  tree is ``level + 1`` and the space needed is the summation of the number of
+  blocks on each level.
+
+- For inode-rooted btrees where the records in the top level do not fit in the
+  inode fork area, the height is ``level + 2``, the space needed is the
+  summation of the number of blocks on each level, and the inode fork points to
+  the root block.
+
+- For inode-rooted btrees where the records in the top level can be stored in
+  the inode fork area, then the root block can be stored in the inode, the
+  height is ``level + 1``, and the space needed is one less than the summation
+  of the number of blocks on each level.
+  This only becomes relevant when non-bmap btrees gain the ability to root in
+  an inode, which is a future patchset and only included here for completeness.
+
+.. _newbt:
+
+Reserving New B+Tree Blocks
+```````````````````````````
+
+Once repair knows the number of blocks needed for the new btree, it allocates
+those blocks using the free space information.
+Each reserved extent is tracked separately by the btree builder state data.
+To improve crash resilience, the reservation code also logs an Extent Freeing
+Intent (EFI) item in the same transaction as each space allocation and attaches
+its in-memory ``struct xfs_extent_free_item`` object to the space reservation.
+If the system goes down, log recovery will use the unfinished EFIs to free the
+unused space, the free space, leaving the filesystem unchanged.
+
+Each time the btree builder claims a block for the btree from a reserved
+extent, it updates the in-memory reservation to reflect the claimed space.
+Block reservation tries to allocate as much contiguous space as possible to
+reduce the number of EFIs in play.
+
+While repair is writing these new btree blocks, the EFIs created for the space
+reservations pin the tail of the ondisk log.
+It's possible that other parts of the system will remain busy and push the head
+of the log towards the pinned tail.
+To avoid livelocking the filesystem, the EFIs must not pin the tail of the log
+for too long.
+To alleviate this problem, the dynamic relogging capability of the deferred ops
+mechanism is reused here to commit a transaction at the log head containing an
+EFD for the old EFI and new EFI at the head.
+This enables the log to release the old EFI to keep the log moving forwards.
+
+EFIs have a role to play during the commit and reaping phases; please see the
+next section and the section about :ref:`reaping<reaping>` for more details.
+
+Proposed patchsets are the
+`bitmap rework
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-bitmap-rework>`_
+and the
+`preparation for bulk loading btrees
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-prep-for-bulk-loading>`_.
+
+
+Writing the New Tree
+````````````````````
+
+This part is pretty simple -- the btree builder (``xfs_btree_bulkload``) claims
+a block from the reserved list, writes the new btree block header, fills the
+rest of the block with records, and adds the new leaf block to a list of
+written blocks.
+Sibling pointers are set every time a new block is added to the level.
+When it finishes writing the record leaf blocks, it moves on to the node
+blocks.
+To fill a node block, it walks each block in the next level down in the tree
+to compute the relevant keys and write them into the parent node.
+When it reaches the root level, it is ready to commit the new btree!
+
+The first step to commit the new btree is to persist the btree blocks to disk
+synchronously.
+This is a little complicated because a new btree block could have been freed
+in the recent past, so the builder must use ``xfs_buf_delwri_queue_here`` to
+remove the (stale) buffer from the AIL list before it can write the new blocks
+to disk.
+Blocks are queued for IO using a delwri list and written in one large batch
+with ``xfs_buf_delwri_submit``.
+
+Once the new blocks have been persisted to disk, control returns to the
+individual repair function that called the bulk loader.
+The repair function must log the location of the new root in a transaction,
+clean up the space reservations that were made for the new btree, and reap the
+old metadata blocks:
+
+1. Commit the location of the new btree root.
+
+2. For each incore reservation:
+
+   a. Log Extent Freeing Done (EFD) items for all the space that was consumed
+      by the btree builder.  The new EFDs must point to the EFIs attached to
+      the reservation to prevent log recovery from freeing the new blocks.
+
+   b. For unclaimed portions of incore reservations, create a regular deferred
+      extent free work item to be free the unused space later in the
+      transaction chain.
+
+   c. The EFDs and EFIs logged in steps 2a and 2b must not overrun the
+      reservation of the committing transaction.
+      If the btree loading code suspects this might be about to happen, it must
+      call ``xrep_defer_finish`` to clear out the deferred work and obtain a
+      fresh transaction.
+
+3. Clear out the deferred work a second time to finish the commit and clean
+   the repair transaction.
+
+The transaction rolling in steps 2c and 3 represent a weakness in the repair
+algorithm, because a log flush and a crash before the end of the reap step can
+result in space leaking.
+Online repair functions minimize the chances of this occuring by using very
+large transactions, which each can accomodate many thousands of block freeing
+instructions.
+Repair moves on to reaping the old blocks, which will be presented in a
+subsequent :ref:`section<reaping>` after a few case studies of bulk loading.
+
+Case Study: Rebuilding the Inode Index
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild the inode index btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_inobt_rec``
+   records from the inode chunk information and a bitmap of the old inode btree
+   blocks.
+
+2. Append the records to an xfarray in inode order.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the inode btree.
+   If the free space inode btree is enabled, call it again to estimate the
+   geometry of the finobt.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+   If the free space inode btree is enabled, call it again to load the finobt.
+
+6. Commit the location of the new btree root block(s) to the AGI.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows.
+
+The inode btree maps inumbers to the ondisk location of the associated
+inode records, which means that the inode btrees can be rebuilt from the
+reverse mapping information.
+Reverse mapping records with an owner of ``XFS_RMAP_OWN_INOBT`` marks the
+location of the old inode btree blocks.
+Each reverse mapping record with an owner of ``XFS_RMAP_OWN_INODES`` marks the
+location of at least one inode cluster buffer.
+A cluster is the smallest number of ondisk inodes that can be allocated or
+freed in a single transaction; it is never smaller than 1 fs block or 4 inodes.
+
+For the space represented by each inode cluster, ensure that there are no
+records in the free space btrees nor any records in the reference count btree.
+If there are, the space metadata inconsistencies are reason enough to abort the
+operation.
+Otherwise, read each cluster buffer to check that its contents appear to be
+ondisk inodes and to decide if the file is allocated
+(``xfs_dinode.i_mode != 0``) or free (``xfs_dinode.i_mode == 0``).
+Accumulate the results of successive inode cluster buffer reads until there is
+enough information to fill a single inode chunk record, which is 64 consecutive
+numbers in the inumber keyspace.
+If the chunk is sparse, the chunk record may include holes.
+
+Once the repair function accumulates one chunk's worth of data, it calls
+``xfarray_append`` to add the inode btree record to the xfarray.
+This xfarray is walked twice during the btree creation step -- once to populate
+the inode btree with all inode chunk records, and a second time to populate the
+free inode btree with records for chunks that have free non-sparse inodes.
+The number of records for the inode btree is the number of xfarray records,
+but the record count for the free inode btree has to be computed as inode chunk
+records are stored in the xfarray.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding the Space Reference Counts
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild the reference count btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_refcount_irec``
+   records for any space having more than one reverse mapping and add them to
+   the xfarray.
+   Any records owned by ``XFS_RMAP_OWN_COW`` are also added to the xfarray.
+   Use any records owned by ``XFS_RMAP_OWN_REFC`` to create a bitmap of old
+   refcount btree blocks.
+
+2. Sort the records in physical extent order, putting the CoW staging extents
+   at the end of the xfarray.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+6. Commit the location of new btree root block to the AGF.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows; the same algorithm is used by ``xfs_repair`` to
+generate refcount information from reverse mapping records.
+
+Reverse mapping records are used to rebuild the reference count information.
+Reference counts are required for correct operation of copy on write for shared
+file data.
+Imagine the reverse mapping entries as rectangles representing extents of
+physical blocks, and that the rectangles can be laid down to allow them to
+overlap each other.
+From the diagram below, it is apparent that a reference count record must start
+or end wherever the height of the stack changes.
+In other words, the record emission stimulus is level-triggered::
+
+                        █    ███
+              ██      █████ ████   ███        ██████
+        ██   ████     ███████████ ████     █████████
+        ████████████████████████████████ ███████████
+        ^ ^  ^^ ^^    ^ ^^ ^^^  ^^^^  ^ ^^ ^  ^     ^
+        2 1  23 21    3 43 234  2123  1 01 2  3     0
+
+The ondisk reference count btree does not store the refcount == 0 cases because
+the free space btree already records which blocks are free.
+Extents being used to stage copy-on-write operations should be the only records
+with refcount == 1.
+Single-owner file blocks aren't recorded in either the free space or the
+reference count btrees.
+
+Given the reverse mapping btree which orders records by physical block number,
+a starting physical block (``sp``), a bag-like data structure to hold mappings
+that cover ``sp``, and the next physical block where the level changes
+(``np``), reference count information is constructed from reverse mapping data
+as follows:
+
+While there are still unprocessed mappings in the reverse mapping btree:
+
+1. Set ``sp`` to the physical block of the next unprocessed reverse mapping
+   record.
+
+2. Add to the bag all the reverse mappings where ``rm_startblock`` == ``sp``.
+
+3. Set ``np`` to the physical block where the bag size will change.
+   This is the minimum of (``rm_startblock`` of the next unprocessed mapping)
+   and (``rm_startblock`` + ``rm_blockcount`` of each mapping in the bag).
+
+4. Record the bag size as ``old_bag_size``.
+
+5. While the bag isn't empty,
+
+   a. Remove from the bag all mappings where ``rm_startblock`` +
+      ``rm_blockcount`` == ``np``.
+
+   b. Add to the bag all reverse mappings where ``rm_startblock`` == ``np``.
+
+   c. If the bag size isn't ``old_bag_size``, store the refcount record
+      ``(sp, np - sp, old_bag_size)`` in the refcount xfarray.
+
+   d. If the bag is empty, break out of this inner loop.
+
+   e. Set ``old_bag_size`` to ``bag_size``.
+
+   f. Set ``sp`` = ``np``.
+
+   g. Set ``np`` to the physical block where the bag size will change.
+      Go to step 3 above.
+
+The bag-like structure in this case is a type 2 xfarray as discussed in the
+:ref:`xfarray access patterns<xfarray_access_patterns>` section.
+Reverse mappings are added to the bag using ``xfarray_store_anywhere`` and
+removed via ``xfarray_unset``.
+Bag members are examined through ``xfarray_iter`` loops.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding File Fork Mapping Indices
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild a data/attr fork mapping btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_bmbt_rec``
+   records from the reverse mapping records for that inode and fork.
+   Append these records to an xfarray.
+   Compute the bitmap of the old bmap btree blocks from the ``BMBT_BLOCK``
+   records.
+
+2. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+3. Sort the records in file offset order.
+
+4. If the extent records would fit in the inode fork immediate area, commit the
+   records to that immediate area and skip to step 8.
+
+5. Allocate the number of blocks computed in the previous step.
+
+6. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+7. Commit the new btree root block to the inode fork immediate area.
+
+8. Reap the old btree blocks using the bitmap created in step 1.
+
+There are some complications here:
+First, it's possible to move the fork offset to adjust the sizes of the
+immediate areas if the data and attr forks are not both in BMBT format.
+Second, if there are sufficiently few fork mappings, it may be possible to use
+EXTENTS format instead of BMBT, which may require a conversion.
+Third, the incore extent map must be reloaded carefully to avoid disturbing
+any delayed allocation extents.
+
+The proposed patchset is the
+`file repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-inodes>`_
+series.
+
+.. _reaping:
+
+Reaping Old Metadata Blocks
+---------------------------
+
+Whenever online fsck builds a new data structure to replace one that is
+suspect, there is a question of how to find and dispose of the blocks that
+belonged to the old structure.
+The laziest method of course is not to deal with them at all, but this slowly
+leads to service degradations as space leaks out of the filesystem.
+Hopefully, someone will schedule a rebuild of the free space information to
+plug all those leaks.
+Offline repair rebuilds all space metadata after recording the usage of
+the files and directories that it decides not to clear, hence it can build new
+structures in the discovered free space and avoid the question of reaping.
+
+As part of a repair, online fsck relies heavily on the reverse mapping records
+to find space that is owned by the corresponding rmap owner yet truly free.
+Cross referencing rmap records with other rmap records is necessary because
+there may be other data structures that also think they own some of those
+blocks (e.g. crosslinked trees).
+Permitting the block allocator to hand them out again will not push the system
+towards consistency.
+
+For space metadata, the process of finding extents to dispose of generally
+follows this format:
+
+1. Create a bitmap of space used by data structures that must be preserved.
+   The space reservations used to create the new metadata can be used here if
+   the same rmap owner code is used to denote all of the objects being rebuilt.
+
+2. Survey the reverse mapping data to create a bitmap of space owned by the
+   same ``XFS_RMAP_OWN_*`` number for the metadata that is being preserved.
+
+3. Use the bitmap disunion operator to subtract (1) from (2).
+   The remaining set bits represent candidate extents that could be freed.
+   The process moves on to step 4 below.
+
+Repairs for file-based metadata such as extended attributes, directories,
+symbolic links, quota files and realtime bitmaps are performed by building a
+new structure attached to a temporary file and swapping the forks.
+Afterward, the mappings in the old file fork are the candidate blocks for
+disposal.
+
+The process for disposing of old extents is as follows:
+
+4. For each candidate extent, count the number of reverse mapping records for
+   the first block in that extent that do not have the same rmap owner for the
+   data structure being repaired.
+
+   - If zero, the block has a single owner and can be freed.
+
+   - If not, the block is part of a crosslinked structure and must not be
+     freed.
+
+5. Starting with the next block in the extent, figure out how many more blocks
+   have the same zero/nonzero other owner status as that first block.
+
+6. If the region is crosslinked, delete the reverse mapping entry for the
+   structure being repaired and move on to the next region.
+
+7. If the region is to be freed, mark any corresponding buffers in the buffer
+   cache as stale to prevent log writeback.
+
+8. Free the region and move on.
+
+However, there is one complication to this procedure.
+Transactions are of finite size, so the reaping process must be careful to roll
+the transactions to avoid overruns.
+Overruns come from two sources:
+
+a. EFIs logged on behalf of space that is no longer occupied
+
+b. Log items for buffer invalidations
+
+This is also a window in which a crash during the reaping process can leak
+blocks.
+As stated earlier, online repair functions use very large transactions to
+minimize the chances of this occurring.
+
+The proposed patchset is the
+`preparation for bulk loading btrees
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-prep-for-bulk-loading>`_
+series.
+
+Case Study: Reaping After a Regular Btree Repair
+````````````````````````````````````````````````
+
+Old reference count and inode btrees are the easiest to reap because they have
+rmap records with special owner codes: ``XFS_RMAP_OWN_REFC`` for the refcount
+btree, and ``XFS_RMAP_OWN_INOBT`` for the inode and free inode btrees.
+Creating a list of extents to reap the old btree blocks is quite simple,
+conceptually:
+
+1. Lock the relevant AGI/AGF header buffers to prevent allocation and frees.
+
+2. For each reverse mapping record with an rmap owner corresponding to the
+   metadata structure being rebuilt, set the corresponding range in a bitmap.
+
+3. Walk the current data structures that have the same rmap owner.
+   For each block visited, clear that range in the above bitmap.
+
+4. Each set bit in the bitmap represents a block that could be a block from the
+   old data structures and hence is a candidate for reaping.
+   In other words, ``(rmap_records_owned_by & ~blocks_reachable_by_walk)``
+   are the blocks that might be freeable.
+
+If it is possible to maintain the AGF lock throughout the repair (which is the
+common case), then step 2 can be performed at the same time as the reverse
+mapping record walk that creates the records for the new btree.
+
+Case Study: Rebuilding the Free Space Indices
+`````````````````````````````````````````````
+
+The high level process to rebuild the free space indices is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_alloc_rec_incore``
+   records from the gaps in the reverse mapping btree.
+
+2. Append the records to an xfarray.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for each new tree.
+
+4. Allocate the number of blocks computed in the previous step from the free
+   space information collected.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks for the free space by block index.
+   Call it again for the free space by length index.
+
+6. Commit the locations of the new btree root blocks to the AGF.
+
+7. Reap the old btree blocks by looking for space that is not recorded by the
+   reverse mapping btree, the new free space btrees, or the AGFL.
+
+Repairing the free space btrees has three key complications over a regular
+btree repair:
+
+First, free space is not explicitly tracked in the reverse mapping records.
+Hence, the new free space records must be inferred from gaps in the physical
+space component of the keyspace of the reverse mapping btree.
+
+Second, free space repairs cannot use the common btree reservation code because
+new blocks are reserved out of the free space btrees.
+This is impossible when repairing the free space btrees themselves.
+However, repair holds the AGF buffer lock for the duration of the free space
+index reconstruction, so it can use the collected free space information to
+supply the blocks for the new free space btrees.
+It is not necessary to back each reserved extent with an EFI because the new
+free space btrees are constructed in what the ondisk filesystem thinks is
+unowned space.
+However, if reserving blocks for the new btrees from the collected free space
+information changes the number of free space records, repair must re-estimate
+the new free space btree geometry with the new record count until the
+reservation is sufficient.
+As part of committing the new btrees, repair must ensure that reverse mappings
+are created for the reserved blocks and that unused reserved blocks are
+inserted into the free space btrees.
+Deferrred rmap and freeing operations are used to ensure that this transition
+is atomic, similar to the other btree repair functions.
+
+Third, finding the blocks to reap after the repair is not overly
+straightforward.
+Blocks for the free space btrees and the reverse mapping btrees are supplied by
+the AGFL.
+Blocks put onto the AGFL have reverse mapping records with the owner
+``XFS_RMAP_OWN_AG``.
+This ownership is retained when blocks move from the AGFL into the free space
+btrees or the reverse mapping btrees.
+When repair walks reverse mapping records to synthesize free space records, it
+creates a bitmap (``ag_owner_bitmap``) of all the space claimed by
+``XFS_RMAP_OWN_AG`` records.
+The repair context maintains a second bitmap corresponding to the rmap btree
+blocks and the AGFL blocks (``rmap_agfl_bitmap``).
+When the walk is complete, the bitmap disunion operation ``(ag_owner_bitmap &
+~rmap_agfl_bitmap)`` computes the extents that are used by the old free space
+btrees.
+These blocks can then be reaped using the methods outlined above.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+.. _rmap_reap:
+
+Case Study: Reaping After Repairing Reverse Mapping Btrees
+``````````````````````````````````````````````````````````
+
+Old reverse mapping btrees are less difficult to reap after a repair.
+As mentioned in the previous section, blocks on the AGFL, the two free space
+btree blocks, and the reverse mapping btree blocks all have reverse mapping
+records with ``XFS_RMAP_OWN_AG`` as the owner.
+The full process of gathering reverse mapping records and building a new btree
+are described in the case study of
+:ref:`live rebuilds of rmap data <rmap_repair>`, but a crucial point from that
+discussion is that the new rmap btree will not contain any records for the old
+rmap btree, nor will the old btree blocks be tracked in the free space btrees.
+The list of candidate reaping blocks is computed by setting the bits
+corresponding to the gaps in the new rmap btree records, and then clearing the
+bits corresponding to extents in the free space btrees and the current AGFL
+blocks.
+The result ``(new_rmapbt_gaps & ~(agfl | bnobt_records))`` are reaped using the
+methods outlined above.
+
+The rest of the process of rebuildng the reverse mapping btree is discussed
+in a separate :ref:`case study<rmap_repair>`.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding the AGFL
+```````````````````````````````
+
+The allocation group free block list (AGFL) is repaired as follows:
+
+1. Create a bitmap for all the space that the reverse mapping data claims is
+   owned by ``XFS_RMAP_OWN_AG``.
+
+2. Subtract the space used by the two free space btrees and the rmap btree.
+
+3. Subtract any space that the reverse mapping data claims is owned by any
+   other owner, to avoid re-adding crosslinked blocks to the AGFL.
+
+4. Once the AGFL is full, reap any blocks leftover.
+
+5. The next operation to fix the freelist will right-size the list.

[PATCH 09/14] xfs: document online file metadata repair code

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Add to the fifth chapter of the online fsck design documentation, where
we discuss the details of the data structures and algorithms used by the
kernel to repair file metadata.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  150 ++++++++++++++++++++
 1 file changed, 150 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index c41f089549a0..10709dc74dcb 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -2872,3 +2872,153 @@ The allocation group free block list (AGFL) is repaired as follows:
 4. Once the AGFL is full, reap any blocks leftover.
 
 5. The next operation to fix the freelist will right-size the list.
+
+Inode Record Repairs
+--------------------
+
+Inode records must be handled carefully, because they have both ondisk records
+("dinodes") and an in-memory ("cached") representation.
+There is a very high potential for cache coherency issues if online fsck is not
+careful to access the ondisk metadata *only* when the ondisk metadata is so
+badly damaged that the filesystem cannot load the in-memory representation.
+When online fsck wants to open a damaged file for scrubbing, it must use
+specialized resource acquisition functions that return either the in-memory
+representation *or* a lock on whichever object is necessary to prevent any
+update to the ondisk location.
+
+The only repairs that should be made to the ondisk inode buffers are whatever
+is necessary to get the in-core structure loaded.
+This means fixing whatever is caught by the inode cluster buffer and inode fork
+verifiers, and retrying the ``iget`` operation.
+If the second ``iget`` fails, the repair has failed.
+
+Once the in-memory representation is loaded, repair can lock the inode and can
+subject it to comprehensive checks, repairs, and optimizations.
+Most inode attributes are easy to check and constrain, or are user-controlled
+arbitrary bit patterns; these are both easy to fix.
+Dealing with the data and attr fork extent counts and the file block counts is
+more complicated, because computing the correct value requires traversing the
+forks, or if that fails, leaving the fields invalid and waiting for the fork
+fsck functions to run.
+
+The proposed patchset is the
+`inode
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-inodes>`_
+repair series.
+
+Quota Record Repairs
+--------------------
+
+Similar to inodes, quota records ("dquots") also have both ondisk records and
+an in-memory representation, and hence are subject to the same cache coherency
+issues.
+Somewhat confusingly, both are known as dquots in the XFS codebase.
+
+The only repairs that should be made to the ondisk quota record buffers are
+whatever is necessary to get the in-core structure loaded.
+Once the in-memory representation is loaded, the only attributes needing
+checking are obviously bad limits and timer values.
+
+Quota usage counters are checked, repaired, and discussed separately in the
+section about :ref:`live quotacheck <quotacheck>`.
+
+The proposed patchset is the
+`quota
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quota>`_
+repair series.
+
+.. _fscounters:
+
+Freezing to Fix Summary Counters
+--------------------------------
+
+Filesystem summary counters track availability of filesystem resources such
+as free blocks, free inodes, and allocated inodes.
+This information could be compiled by walking the free space and inode indexes,
+but this is a slow process, so XFS maintains a copy in the ondisk superblock
+that should reflect the ondisk metadata, at least when the filesystem has been
+unmounted cleanly.
+For performance reasons, XFS also maintains incore copies of those counters,
+which are key to enabling resource reservations for active transactions.
+Writer threads reserve the worst-case quantities of resources from the
+incore counter and give back whatever they don't use at commit time.
+It is therefore only necessary to serialize on the superblock when the
+superblock is being committed to disk.
+
+The lazy superblock counter feature introduced in XFS v5 took this even further
+by training log recovery to recompute the summary counters from the AG headers,
+which eliminated the need for most transactions even to touch the superblock.
+The only time XFS commits the summary counters is at filesystem unmount.
+To reduce contention even further, the incore counter is implemented as a
+percpu counter, which means that each CPU is allocated a batch of blocks from a
+global incore counter and can satisfy small allocations from the local batch.
+
+The high-performance nature of the summary counters makes it difficult for
+online fsck to check them, since there is no way to quiesce a percpu counter
+while the system is running.
+Although online fsck can read the filesystem metadata to compute the correct
+values of the summary counters, there's no way to hold the value of a percpu
+counter stable, so it's quite possible that the counter will be out of date by
+the time the walk is complete.
+Earlier versions of online scrub would return to userspace with an incomplete
+scan flag, but this is not a satisfying outcome for a system administrator.
+For repairs, the in-memory counters must be stabilize while walking the
+filesystem metadata to get an accurate reading and install it in the percpu
+counter.
+
+To satisfy this requirement, online fsck must prevent other programs in the
+system from initiating new writes to the filesystem, it must disable background
+garbage collection threads, and it must wait for existing writer programs to
+exit the kernel.
+Once that has been established, scrub can walk the AG free space indexes, the
+inode btrees, and the realtime bitmap to compute the correct value of all
+four summary counters.
+This is very similar to a filesystem freeze.
+
+The initial implementation used the actual VFS filesystem freeze mechanism to
+quiesce filesystem activity.
+With the filesystem frozen, it is possible to resolve the counter values with
+exact precision, but there are many problems with calling the VFS methods
+directly:
+
+- Other programs can unfreeze the filesystem without our knowledge.
+  This leads to incorrect scan results and incorrect repairs.
+
+- Adding an extra lock to prevent others from thawing the filesystem required
+  the addition of a ``->freeze_super`` function to wrap ``freeze_fs()``.
+  This in turn caused other subtle problems because it turns out that the VFS
+  ``freeze_super`` and ``thaw_super`` functions can drop the last reference to
+  the VFS superblock, and any subsequent access becomes a UAF bug!
+  This can happen if the filesystem is unmounted while the underlying block
+  device has frozen the filesystem.
+  This problem could be solved by grabbing extra references to the superblock,
+  but it felt suboptimal given the other inadequacies of this approach:
+
+- The log need not be quiesced to check the summary counters, but a VFS freeze
+  initiates one anyway.
+  This adds unnecessary runtime to live fscounter fsck operations.
+
+- Quiescing the log means that XFS flushes the (possibly incorrect) counters to
+  disk as part of cleaning the log.
+
+- A bug in the VFS meant that freeze could complete even when sync_filesystem
+  fails to flush the filesystem and returns an error.
+  This bug was fixed in Linux 5.17.
+
+The author established that the only component of online fsck that requires the
+ability to freeze the filesystem is the fscounter scrubber, so the code for
+this could be localized to that source file.
+fscounter freeze behaves the same as the VFS freeze method, except:
+
+- The final freeze state is set one higher than ``SB_FREEZE_COMPLETE`` to
+  prevent other threads from thawing the filesystem.
+
+- It does not quiesce the log.
+
+With this code in place, it is now possible to pause the filesystem for just
+long enough to check and correct the summary counters.
+
+The proposed patchset is the
+`summary counter cleanup
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-fscounters>`_
+series.

[PATCH 10/14] xfs: document full filesystem scans for online fsck

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Certain parts of the online fsck code need to scan every file in the
entire filesystem.  It is not acceptable to block the entire filesystem
while this happens, which means that we need to be clever in allowing
scans to coordinate with ongoing filesystem updates.  We also need to
hook the filesystem so that regular updates propagate to the staging
records.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  677 ++++++++++++++++++++
 1 file changed, 677 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 10709dc74dcb..5ab2d76ad694 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -3022,3 +3022,680 @@ The proposed patchset is the
 `summary counter cleanup
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-fscounters>`_
 series.
+
+Full Filesystem Scans
+---------------------
+
+Certain types of metadata can only be checked by walking every file in the
+entire filesystem to record observations and comparing the observations against
+what's recorded on disk.
+Like every other type of online repair, repairs are made by writing those
+observations to disk in a replacement structure and committing it atomically.
+However, it is not practical to shut down the entire filesystem to examine
+hundreds of billions of files because the downtime would be excessive.
+Therefore, online fsck must build the infrastructure to manage a live scan of
+all the files in the filesystem.
+There are two questions that need to be solved to perform a live walk:
+
+- How does scrub manage the scan while it is collecting data?
+
+- How does the scan keep abreast of changes being made to the system by other
+  threads?
+
+.. _iscan:
+
+Coordinated Inode Scans
+```````````````````````
+
+In the original Unix filesystems of the 1970s, each directory entry contained
+an index number (*inumber*) which was used as an index into on ondisk array
+(*itable*) of fixed-size records (*inodes*) describing a file's attributes and
+its data block mapping.
+This system is described by J. Lions, `"inode (5659)"
+<http://www.lemis.com/grog/Documentation/Lions/>`_ in *Lions' Commentary on
+UNIX, 6th Edition*, (Dept. of Computer Science, the University of New South
+Wales, November 1977), pp. 18-2; and later by D. Ritchie and K. Thompson,
+`"Implementation of the File System"
+<https://archive.org/details/bstj57-6-1905/page/n8/mode/1up>`_, from *The UNIX
+Time-Sharing System*, (The Bell System Technical Journal, July 1978), pp.
+1913-4.
+
+XFS retains most of this design, except now inumbers are search keys over all
+the space in the data section filesystem.
+They form a continuous keyspace that can be expressed as a 64-bit integer,
+though the inodes themselves are sparsely distributed within the keyspace.
+Scans proceed in a linear fashion across the inumber keyspace, starting from
+``0x0`` and ending at ``0xFFFFFFFFFFFFFFFF``.
+Naturally, a scan through a keyspace requires a scan cursor object to track the
+scan progress.
+Because this keyspace is sparse, this cursor contains two parts.
+The first part of this scan cursor object tracks the inode that will be
+examined next; call this the examination cursor.
+Somewhat less obviously, the scan cursor object must also track which parts of
+the keyspace have already been visited, which is critical for deciding if a
+concurrent filesystem update needs to be incorporated into the scan data.
+Call this the visited inode cursor.
+
+Advancing the scan cursor is a multi-step process encapsulated in
+``xchk_iscan_iter``:
+
+1. Lock the AGI buffer of the AG containing the inode pointed to by the visited
+   inode cursor.
+   This guarantee that inodes in this AG cannot be allocated or freed while
+   advancing the cursor.
+
+2. Use the per-AG inode btree to look up the next inumber after the one that
+   was just visited, since it may not be keyspace adjacent.
+
+3. If there are no more inodes left in this AG:
+
+   a. Move the examination cursor to the point of the inumber keyspace that
+      corresponds to the start of the next AG.
+
+   b. Adjust the visited inode cursor to indicate that it has "visited" the
+      last possible inode in the current AG's inode keyspace.
+      XFS inumbers are segmented, so the cursor needs to be marked as having
+      visited the entire keyspace up to just before the start of the next AG's
+      inode keyspace.
+
+   c. Unlock the AGI and return to step 1 if there are unexamined AGs in the
+      filesystem.
+
+   d. If there are no more AGs to examine, set both cursors to the end of the
+      inumber keyspace.
+      The scan is now complete.
+
+4. Otherwise, there is at least one more inode to scan in this AG:
+
+   a. Move the examination cursor ahead to the next inode marked as allocated
+      by the inode btree.
+
+   b. Adjust the visited inode cursor to point to the inode just prior to where
+      the examination cursor is now.
+      Because the scanner holds the AGI buffer lock, no inodes could have been
+      created in the part of the inode keyspace that the visited inode cursor
+      just advanced.
+
+5. Get the incore inode for the inumber of the examination cursor.
+   By maintaining the AGI buffer lock until this point, the scanner knows that
+   it was safe to advance the examination cursor across the entire keyspace,
+   and that it has stabilized this next inode so that it cannot disappear from
+   the filesystem until the scan releases the incore inode.
+
+6. Drop the AGI lock and return the incore inode to the caller.
+
+Online fsck functions scan all files in the filesystem as follows:
+
+1. Start a scan by calling ``xchk_iscan_start``.
+
+2. Advance the scan cursor (``xchk_iscan_iter``) to get the next inode.
+   If one is provided:
+
+   a. Lock the inode to prevent updates during the scan.
+
+   b. Scan the inode.
+
+   c. While still holding the inode lock, adjust the visited inode cursor
+      (``xchk_iscan_mark_visited``) to point to this inode.
+
+   d. Unlock and release the inode.
+
+8. Call ``xchk_iscan_finish`` to complete the scan.
+
+There are subtleties with the inode cache that complicate grabbing the incore
+inode for the caller.
+Obviously, it is an absolute requirement that the inode metadata be consistent
+enough to load it into the inode cache.
+Second, if the incore inode is stuck in some intermediate state, the scan
+coordinator must release the AGI and push the main filesystem to get the inode
+back into a loadable state.
+
+The proposed patches are the
+`inode scanner
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-iscan>`_
+series.
+
+Inode Management
+````````````````
+
+In regular filesystem code, references to allocated XFS incore inodes are
+always obtained (``xfs_iget``) outside of transaction context because the
+creation of the incore context for ane xisting file does not require metadata
+updates.
+However, it is important to note that references to incore inodes obtained as
+part of file creation must be performed in transaction context because the
+filesystem must ensure the atomicity of the ondisk inode btree index updates
+and the initialization of the actual ondisk inode.
+
+References to incore inodes are always released (``xfs_irele``) outside of
+transaction context because there are a handful of activities that might
+require ondisk updates:
+
+- The VFS may decide to kick off writeback as part of a ``DONTCACHE`` inode
+  release.
+
+- Speculative preallocations need to be unreserved.
+
+- An unlinked file may have lost its last reference, in which case the entire
+  file must be inactivated, which involves releasing all of its resources in
+  the ondisk metadata and freeing the inode.
+
+These activities are collectively called inode inactivation.
+Inactivation has two parts -- the VFS part, which initiates writeback on all
+dirty file pages, and the XFS part, which cleans up XFS-specific information
+and frees the inode if it was unlinked.
+If the inode is unlinked (or unconnected after a file handle operation), the
+kernel drops the inode into the inactivation machinery immediately.
+
+During normal operation, resource acquisition for an update follows this order
+to avoid deadlocks:
+
+1. Inode reference (``iget``).
+
+2. Filesystem freeze protection, if repairing (``mnt_want_write_file``).
+
+3. Inode ``IOLOCK`` (VFS ``i_rwsem``) lock to control file IO.
+
+4. Inode ``MMAPLOCK`` (page cache ``invalidate_lock``) lock for operations that
+   can update page cache mappings.
+
+5. Log feature enablement.
+
+6. Transaction log space grant.
+
+7. Space on the data and realtime devices for the transaction.
+
+8. Incore dquot references, if a file is being repaired.
+   Note that they are not locked, merely acquired.
+
+9. Inode ``ILOCK`` for file metadata updates.
+
+10. AG header buffer locks / Realtime metadata inode ILOCK.
+
+11. Realtime metadata buffer locks, if applicable.
+
+12. Extent mapping btree blocks, if applicable.
+
+Resources are often released in the reverse order, though this is not required.
+However, online fsck differs from regular XFS operations because it may examine
+an object that normally is acquired in a later stage of the locking order, and
+then decide to cross-reference the object with an object that is acquired
+earlier in the order.
+The next few sections detail the specific ways in which online fsck takes care
+to avoid deadlocks.
+
+iget and irele During a Scrub
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+An inode scan performed on behalf of a scrub operation runs in transaction
+context, and possibly with resources already locked and bound to it.
+This isn't much of a problem for ``iget`` since it can operate in the context
+of an existing transaction, as long as all of the bound resources are acquired
+before the inode reference in the regular filesystem.
+
+When the VFS ``iput`` function is given a linked inode with no other
+references, it normally puts the inode on an LRU list in the hope that it can
+save time if another process re-opens the file before the system runs out
+of memory and frees it.
+Filesystem callers can short-circuit the LRU process by setting a ``DONTCACHE``
+flag on the inode to cause the kernel to try to drop the inode into the
+inactivation machinery immediately.
+
+In the past, inactivation was always done from the process that dropped the
+inode, which was a problem for scrub because scrub may already hold a
+transaction, and XFS does not support nesting transactions.
+On the other hand, if there is no scrub transaction, it is desirable to drop
+otherwise unused inodes immediately to avoid polluting caches.
+To capture these nuances, the online fsck code has a separate ``xchk_irele``
+function to set or clear the ``DONTCACHE`` flag to get the required release
+behavior.
+
+Proposed patchsets include fixing
+`scrub iget usage
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-iget-fixes>`_ and
+`dir iget usage
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-dir-iget-fixes>`_.
+
+Locking Inodes
+^^^^^^^^^^^^^^
+
+In regular filesystem code, the VFS and XFS will acquire multiple IOLOCK locks
+in a well-known order: parent → child when updating the directory tree, and
+``struct inode`` address order otherwise.
+For regular files, the MMAPLOCK can be acquired after the IOLOCK to stop page
+faults.
+If two MMAPLOCKs must be acquired, they are acquired in ``struct
+address_space`` order.
+Due to the structure of existing filesystem code, IOLOCKs and MMAPLOCKs must be
+acquired before transactions are allocated.
+If two ILOCKs must be acquired, they are acquired in inumber order.
+
+Inode lock acquisition must be done carefully during a coordinated inode scan.
+Online fsck cannot abide these conventions, because for a directory tree
+scanner, the scrub process holds the IOLOCK of the file being scanned and it
+needs to take the IOLOCK of the file at the other end of the directory link.
+If the directory tree is corrupt because it contains a cycle, ``xfs_scrub``
+cannot use the regular inode locking functions and avoid becoming trapped in an
+ABBA deadlock.
+
+Solving both of these problems is straightforward -- any time online fsck
+needs to take a second lock of the same class, it uses trylock to avoid an ABBA
+deadlock.
+If the trylock fails, scrub drops all inode locks and use trylock loops to
+(re)acquire all necessary resources.
+Trylock loops enable scrub to check for pending fatal signals, which is how
+scrub avoids deadlocking the filesystem or becoming an unresponsive process.
+However, trylock loops means that online fsck must be prepared to measure the
+resource being scrubbed before and after the lock cycle to detect changes and
+react accordingly.
+
+.. _dirparent:
+
+Case Study: Finding a Directory Parent
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Consider the directory parent pointer repair code as an example.
+Online fsck must verify that the dotdot dirent of a directory points up to a
+parent directory, and that the parent directory contains exactly one dirent
+pointing down to the child directory.
+Fully validating this relationship (and repairing it if possible) requires a
+walk of every directory on the filesystem while holding the child locked, and
+while updates to the directory tree are being made.
+The coordinated inode scan provides a way to walk the filesystem without the
+possibility of missing an inode.
+The child directory is kept locked to prevent updates to the dotdot dirent, but
+if the scanner fails to lock a parent, it can drop and relock both the child
+and the prospective parent.
+If the dotdot entry changes while the directory is unlocked, then a move or
+rename operation must have changed the child's parentage, and the scan can
+exit early.
+
+The proposed patchset is the
+`directory repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-dirs>`_
+series.
+
+.. _fshooks:
+
+Filesystem Hooks
+`````````````````
+
+The second piece of support that online fsck functions need during a full
+filesystem scan is the ability to stay informed about updates being made by
+other threads in the filesystem, since comparisons against the past are useless
+in a dynamic environment.
+Two pieces of Linux kernel infrastructure enable online fsck to monitor regular
+filesystem operations: filesystem hooks and :ref:`static keys<jump_labels>`.
+
+Filesystem hooks convey information about an ongoing filesystem operation to
+a downstream consumer.
+In this case, the downstream consumer is always an online fsck function.
+Because multiple fsck functions can run in parallel, online fsck uses the Linux
+notifier call chain facility to dispatch updates to any number of interested
+fsck processes.
+Call chains are a dynamic list, which means that they can be configured at
+run time.
+Because these hooks are private to the XFS module, the information passed along
+contains exactly what the checking function needs to update its observations.
+
+The current implementation of XFS hooks uses SRCU notifier chains to reduce the
+impact to highly threaded workloads.
+Regular blocking notifier chains use a rwsem and seem to have a much lower
+overhead for single-threaded applications.
+However, it may turn out that the combination of blocking chains and static
+keys are a more performant combination; more study is needed here.
+
+The following pieces are necessary to hook a certain point in the filesystem:
+
+- A ``struct xfs_hooks`` object must be embedded in a convenient place such as
+  a well-known incore filesystem object.
+
+- Each hook must define an action code and a structure containing more context
+  about the action.
+
+- Hook providers should provide appropriate wrapper functions and structs
+  around the ``xfs_hooks`` and ``xfs_hook`` objects to take advantage of type
+  checking to ensure correct usage.
+
+- A callsite in the regular filesystem code must be chosen to call
+  ``xfs_hooks_call`` with the action code and data structure.
+  This place should be adjacent to (and not earlier than) the place where
+  the filesystem update is committed to the transaction.
+  In general, when the filesystem calls a hook chain, it should be able to
+  handle sleeping and should not be vulnerable to memory reclaim or locking
+  recursion.
+  However, the exact requirements are very dependent on the context of the hook
+  caller and the callee.
+
+- The online fsck function should define a structure to hold scan data, a lock
+  to coordinate access to the scan data, and a ``struct xfs_hook`` object.
+  The scanner function and the regular filesystem code must acquire resources
+  in the same order; see the next section for details.
+
+- The online fsck code must contain a C function to catch the hook action code
+  and data structure.
+  If the object being updated has already been visited by the scan, then the
+  hook information must be applied to the scan data.
+
+- Prior to unlocking inodes to start the scan, online fsck must call
+  ``xfs_hooks_setup`` to initialize the ``struct xfs_hook``, and
+  ``xfs_hooks_add`` to enable the hook.
+
+- Online fsck must call ``xfs_hooks_del`` to disable the hook once the scan is
+  complete.
+
+The number of hooks should be kept to a minimum to reduce complexity.
+Static keys are used to reduce the overhead of filesystem hooks to nearly
+zero when online fsck is not running.
+
+.. _liveupdate:
+
+Live Updates During a Scan
+``````````````````````````
+
+The code paths of the online fsck scanning code and the :ref:`hooked<fshooks>`
+filesystem code look like this::
+
+            other program
+                  ↓
+            inode lock ←────────────────────┐
+                  ↓                         │
+            AG header lock                  │
+                  ↓                         │
+            filesystem function             │
+                  ↓                         │
+            notifier call chain             │    same
+                  ↓                         ├─── inode
+            scrub hook function             │    lock
+                  ↓                         │
+            scan data mutex ←──┐    same    │
+                  ↓            ├─── scan    │
+            update scan data   │    lock    │
+                  ↑            │            │
+            scan data mutex ←──┘            │
+                  ↑                         │
+            inode lock ←────────────────────┘
+                  ↑
+            scrub function
+                  ↑
+            inode scanner
+                  ↑
+            xfs_scrub
+
+These rules must be followed to ensure correct interactions between the
+checking code and the code making an update to the filesystem:
+
+- Prior to invoking the notifier call chain, the filesystem function being
+  hooked must acquire the same lock that the scrub scanning function acquires
+  to scan the inode.
+
+- The scanning function and the scrub hook function must coordinate access to
+  the scan data by acquiring a lock on the scan data.
+
+- Scrub hook function must not add the live update information to the scan
+  observations unless the inode being updated has already been scanned.
+  The scan coordinator has a helper predicate (``xchk_iscan_want_live_update``)
+  for this.
+
+- Scrub hook functions must not change the caller's state, including the
+  transaction that it is running.
+  They must not acquire any resources that might conflict with the filesystem
+  function being hooked.
+
+- The hook function can abort the inode scan to avoid breaking the other rules.
+
+The inode scan APIs are pretty simple:
+
+- ``xchk_iscan_start`` starts a scan
+
+- ``xchk_iscan_iter`` grabs a reference to the next inode in the scan or
+  returns zero if there is nothing left to scan
+
+- ``xchk_iscan_want_live_update`` to decide if an inode has already been
+  visited in the scan.
+  This is critical for hook functions to decide if they need to update the
+  in-memory scan information.
+
+- ``xchk_iscan_mark_visited`` to mark an inode as having been visited in the
+  scan
+
+- ``xchk_iscan_finish`` to finish the scan
+
+The proposed patches are at the start of the
+`online quotacheck
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quota>`_
+series.
+
+.. _quotacheck:
+
+Case Study: Quota Counter Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+It is useful to compare the mount time quotacheck code to the online repair
+quotacheck code.
+Mount time quotacheck does not have to contend with concurrent operations, so
+it does the following:
+
+1. Make sure the ondisk dquots are in good enough shape that all the incore
+   dquots will actually load, and zero the resource usage counters in the
+   ondisk buffer.
+
+2. Walk every inode in the filesystem.
+   Add each file's resource usage to the incore dquot.
+
+3. Walk each incore dquot.
+   If the incore dquot is not being flushed, add the ondisk buffer backing the
+   incore dquot to a delayed write (delwri) list.
+
+4. Write the buffer list to disk.
+
+Like most online fsck functions, online quotacheck can't write to regular
+filesystem objects until the newly collected metadata reflect all filesystem
+state.
+Therefore, online quotacheck records file resource usage to a shadow dquot
+index implemented with a sparse ``xfarray``, and only writes to the real dquots
+once the scan is complete.
+Handling transactional updates is tricky because quota resource usage updates
+are handled in phases to minimize contention on dquots:
+
+1. The inodes involved are joined and locked to a transaction.
+
+2. For each dquot attached to the file:
+
+   a. The dquot is locked.
+
+   b. A quota reservation is added to the dquot's resource usage.
+      The reservation is recorded in the transaction.
+
+   c. The dquot is unlocked.
+
+3. Changes in actual quota usage are tracked in the transaction.
+
+4. At transaction commit time, each dquot is examined again:
+
+   a. The dquot is locked again.
+
+   b. Quota usage changes are logged and unused reservation is given back to
+      the dquot.
+
+   c. The dquot is unlocked.
+
+For online quotacheck, hooks are placed in steps 2 and 4.
+The step 2 hook creates a shadow version of the transaction dquot context
+(``dqtrx``) that operates in a similar manner to the regular code.
+The step 4 hook commits the shadow ``dqtrx`` changes to the shadow dquots.
+Notice that both hooks are called with the inode locked, which is how the
+live update coordinates with the inode scanner.
+
+The quotacheck scan looks like this:
+
+1. Set up a coordinated inode scan.
+
+2. For each inode returned by the inode scan iterator:
+
+   a. Grab and lock the inode.
+
+   b. Determine that inode's resource usage (data blocks, inode counts,
+      realtime blocks) and add that to the shadow dquots for the user, group,
+      and project ids associated with the inode.
+
+   c. Unlock and release the inode.
+
+3. For each dquot in the system:
+
+   a. Grab and lock the dquot.
+
+   b. Check the dquot against the shadow dquots created by the scan and updated
+      by the live hooks.
+
+Live updates are key to being able to walk every quota record without
+needing to hold any locks for a long duration.
+If repairs are desired, the real and shadow dquots are locked and their
+resource counts are set to the values in the shadow dquot.
+
+The proposed patchset is the
+`online quotacheck
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quota>`_
+series.
+
+.. _nlinks:
+
+Case Study: File Link Count Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+File link count checking also uses live update hooks.
+The coordinated inode scanner is used to visit all directories on the
+filesystem, and per-file link count records are stored in a sparse ``xfarray``
+indexed by inumber.
+During the scanning phase, each entry in a directory generates observation
+data as follows:
+
+1. If the entry is a dotdot (``'..'``) entry of the root directory, the
+   directory's parent link count is bumped because the root directory's dotdot
+   entry is self referential.
+
+2. If the entry is a dotdot entry of a subdirectory, the parent's backref
+   count is bumped.
+
+3. If the entry is neither a dot nor a dotdot entry, the target file's parent
+   count is bumped.
+
+4. If the target is a subdirectory, the parent's child link count is bumped.
+
+A crucial point to understand about how the link count inode scanner interacts
+with the live update hooks is that the scan cursor tracks which *parent*
+directories have been scanned.
+In other words, the live updates ignore any update about ``A → B`` when A has
+not been scanned, even if B has been scanned.
+Furthermore, a subdirectory A with a dotdot entry pointing back to B is
+accounted as a backref counter in the shadow data for A, since child dotdot
+entries affect the parent's link count.
+Live update hooks are carefully placed in all parts of the filesystem that
+create, change, or remove directory entries, since those operations involve
+bumplink and droplink.
+
+For any file, the correct link count is the number of parents plus the number
+of child subdirectories.
+Non-directories never have children of any kind.
+The backref information is used to detect inconsistencies in the number of
+links pointing to child subdirectories and the number of dotdot entries
+pointing back.
+
+After the scan completes, the link count of each file can be checked by locking
+both the inode and the shadow data, and comparing the link counts.
+A second coordinated inode scan cursor is used for comparisons.
+Live updates are key to being able to walk every inode without needing to hold
+any locks between inodes.
+If repairs are desired, the inode's link count is set to the value in the
+shadow information.
+If no parents are found, the file must be :ref:`reparented <orphanage>` to the
+orphanage to prevent the file from being lost forever.
+
+The proposed patchset is the
+`file link count repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-nlinks>`_
+series.
+
+.. _rmap_repair:
+
+Case Study: Rebuilding Reverse Mapping Records
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most repair functions follow the same pattern: lock filesystem resources,
+walk the surviving ondisk metadata looking for replacement metadata records,
+and use an :ref:`in-memory array <xfarray>` to store the gathered observations.
+The primary advantage of this approach is the simplicity and modularity of the
+repair code -- code and data are entirely contained within the scrub module,
+do not require hooks in the main filesystem, and are usually the most efficient
+in memory use.
+A secondary advantage of this repair approach is atomicity -- once the kernel
+decides a structure is corrupt, no other threads can access the metadata until
+the kernel finishes repairing and revalidating the metadata.
+
+For repairs going on within a shard of the filesystem, these advantages
+outweigh the delays inherent in locking the shard while repairing parts of the
+shard.
+Unfortunately, repairs to the reverse mapping btree cannot use the "standard"
+btree repair strategy because it must scan every space mapping of every fork of
+every file in the filesystem, and the filesystem cannot stop.
+Therefore, rmap repair foregoes atomicity between scrub and repair.
+It combines a :ref:`coordinated inode scanner <iscan>`, :ref:`live update hooks
+<liveupdate>`, and an :ref:`in-memory rmap btree <xfbtree>` to complete the
+scan for reverse mapping records.
+
+1. Set up an xfbtree to stage rmap records.
+
+2. While holding the locks on the AGI and AGF buffers acquired during the
+   scrub, generate reverse mappings for all AG metadata: inodes, btrees, CoW
+   staging extents, and the internal log.
+
+3. Set up an inode scanner.
+
+4. Hook into rmap updates for the AG being repaired so that the live scan data
+   can receive updates to the rmap btree from the rest of the filesystem during
+   the file scan.
+
+5. For each space mapping found in either fork of each file scanned,
+   decide if the mapping matches the AG of interest.
+   If so:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Use the rmap code to add the record to the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+
+6. For each live update received via the hook, decide if the owner has already
+   been scanned.
+   If so, apply the live update into the scan data:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Replay the operation into the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+      This is performed with an empty transaction to avoid changing the
+      caller's state.
+
+7. When the inode scan finishes, create a new scrub transaction and relock the
+   two AG headers.
+
+8. Compute the new btree geometry using the number of rmap records in the
+   shadow btree, like all other btree rebuilding functions.
+
+9. Allocate the number of blocks computed in the previous step.
+
+10. Perform the usual btree bulk loading and commit to install the new rmap
+    btree.
+
+11. Reap the old rmap btree blocks as discussed in the case study about how
+    to :ref:`reap after rmap btree repair <rmap_reap>`.
+
+12. Free the xfbtree now that it not needed.
+
+The proposed patchset is the
+`rmap repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rmap-btree>`_
+series.

[PATCH 11/14] xfs: document metadata file repair

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

File-based metadata (such as xattrs and directories) can be extremely
large.  To reduce the memory requirements and maximize code reuse, it is
very convenient to create a temporary file, use the regular dir/attr
code to store salvaged information, and then atomically swap the extents
between the file being repaired and the temporary file.  Record the high
level concepts behind how temporary files and atomic content swapping
should work, and then present some case studies of what the actual
repair functions do.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  574 ++++++++++++++++++++
 1 file changed, 574 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 5ab2d76ad694..6cdec62e3f23 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -3256,6 +3256,8 @@ Proposed patchsets include fixing
 `dir iget usage
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-dir-iget-fixes>`_.
 
+.. _ilocking:
+
 Locking Inodes
 ^^^^^^^^^^^^^^
 
@@ -3699,3 +3701,575 @@ The proposed patchset is the
 `rmap repair
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rmap-btree>`_
 series.
+
+Staging Repairs with Temporary Files on Disk
+--------------------------------------------
+
+XFS stores a substantial amount of metadata in file forks: directories,
+extended attributes, symbolic link targets, free space bitmaps and summary
+information for the realtime volume, and quota records.
+File forks map 64-bit logical file fork space extents to physical storage space
+extents, similar to how a memory management unit maps 64-bit virtual addresses
+to physical memory addresses.
+Therefore, file-based tree structures (such as directories and extended
+attributes) use blocks mapped in the file fork offset address space that point
+to other blocks mapped within that same address space, and file-based linear
+structures (such as bitmaps and quota records) compute array element offsets in
+the file fork offset address space.
+
+In the initial iteration of file metadata repair, the damaged metadata blocks
+would be scanned for salvageable data; the extents in the file fork would be
+reaped; and then a new structure would be built in its place.
+This strategy did not survive the introduction of the atomic repair requirement
+expressed earlier in this document.
+The second iteration explored building a second structure at a high offset
+in the fork from the salvage data, reaping the old extents, and using a
+``COLLAPSE_RANGE`` operation to slide the new extents into place.
+This had many drawbacks:
+
+- Array structures are linearly addressed, and the regular filesystem codebase
+  does not have the concept of a linear offset that could be applied to the
+  record offset computation to build an alternate copy.
+
+- Extended attributes are allowed to use the entire attr fork offset address
+  space.
+
+- Even if repair could build an alternate copy of a data structure in a
+  different part of the fork address space, the atomic repair commit
+  requirement means that online repair would have to be able to perform a log
+  assisted ``COLLAPSE_RANGE`` operation to ensure that the old structure was
+  completely replaced.
+
+- A crash after construction of the secondary tree but before the range
+  collapse would leave unreachable blocks in the file fork.
+  This would likely confuse things further.
+
+- Reaping blocks after a repair is not a simple operation, and initiating a
+  reap operation from a restarted range collapse operation during log recovery
+  is daunting.
+
+- Directory entry blocks and quota records record the file fork offset in the
+  header area of each block.
+  An atomic range collapse operation would have to rewrite this part of each
+  block header.
+  Rewriting a single field in block headers is not a huge problem, but it's
+  something to be aware of.
+
+- Each block in a directory or extended attributes btree index contains sibling
+  and child block pointers.
+  Were the atomic commit to use a range collapse operation, each block would
+  have to be rewritten very carefully to preserve the graph structure.
+  Doing this as part of a range collapse means rewriting a large number of
+  blocks repeatedly, which is not conducive to quick repairs.
+
+The third iteration of the design for file metadata repair went for a totally
+new strategy -- create a temporary file in the XFS filesystem, write a new
+structure at the correct offsets into the temporary file, and atomically swap
+the fork mappings (and hence the fork contents) to commit the repair.
+Once the repair is complete, the old fork can be reaped as necessary; if the
+system goes down during the reap, the iunlink code will delete the blocks
+during log recovery.
+
+**Note**: All space usage and inode indices in the filesystem *must* be
+consistent to use a temporary file safely!
+This dependency is the reason why online repair can only use pageable kernel
+memory to stage ondisk space usage information.
+
+Swapping extents with a temporary file still requires a rewrite of the owner
+field of the block headers, but this is *much* simpler than moving tree blocks
+individually.
+Furthermore, the buffer verifiers do not verify owner fields (since they are
+not aware of the inode that owns the block), which makes reaping of old file
+blocks much simpler.
+Extent swapping requires that AG space metadata and the file fork metadata of
+the file being repaired are all consistent with respect to each other, but
+that's already a requirement for correct operation of files in general.
+There is, however, a slight downside -- if the system crashes during the reap
+phase and the fork extents are crosslinked, the iunlink processing will fail
+because freeing space will find the extra reverse mappings and abort.
+
+Temporary files created for repair are similar to ``O_TMPFILE`` files created
+by userspace.
+They are not linked into a directory and the entire file will be reaped when
+the last reference to the file is lost.
+The key differences are that these files must have no access permission outside
+the kernel at all, they must be specially marked to prevent them from being
+opened by handle, and they must never be linked into the directory tree.
+
+Using a Temporary File
+``````````````````````
+
+Online repair code should use the ``xrep_tempfile_create`` function to create a
+temporary file inside the filesystem.
+This allocates an inode, marks the in-core inode private, and attaches it to
+the scrub context.
+These files are hidden from userspace, may not be added to the directory tree,
+and must be kept private.
+
+Temporary files only use two inode locks: the IOLOCK and the ILOCK.
+The MMAPLOCK is not needed here, because there must not be page faults from
+userspace for data fork blocks.
+The usage patterns of these two locks are the same as for any other XFS file --
+access to file data are controlled via the IOLOCK, and access to file metadata
+are controlled via the ILOCK.
+Locking helpers are provided so that the temporary file and its lock state can
+be cleaned up by the scrub context.
+To comply with the nested locking strategy laid out in the :ref:`inode
+locking<ilocking>` section, it is recommended that scrub functions use the
+xrep_tempfile_ilock*_nowait lock helpers.
+
+Data can be written to a temporary file by two means:
+
+1. ``xrep_tempfile_copyin_xfile`` can be used to set the contents of a regular
+   temporary file from an xfile.
+
+2. The regular directory, symbolic link, and extended attribute functions can
+   be used to write to the temporary file.
+
+Once a good copy of a data file has been constructed in a temporary file, it
+must be conveyed to the file being repaired, which is the topic of the next
+section.
+
+The proposed patches are in the
+`realtime summary repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rtsummary>`_
+series.
+
+Atomic Extent Swapping
+----------------------
+
+Once repair builds a temporary file with a new data structure written into
+it, it must commit the new changes into the existing file.
+It is not possible to swap the inumbers of two files, so instead the new
+metadata must replace the old.
+This suggests the need for the ability to swap extents, but the existing extent
+swapping code used by the file defragmenting tool ``xfs_fsr`` is not sufficient
+for online repair because:
+
+a. When the reverse-mapping btree is enabled, the swap code must keep the
+   reverse mapping information up to date with every exchange of mappings.
+   Therefore, it can only exchange one mapping per transaction, and each
+   transaction is independent.
+
+b. Reverse-mapping is critical for the operation of online fsck, so the old
+   defragmentation code (which swapped entire extent forks in a single
+   operation) is not useful here.
+
+c. Defragmentation is assumed to occur between two files with identical
+   contents.
+   For this use case, an incomplete exchange will not result in a user-visible
+   change in file contents, even if the operation is interrupted.
+
+d. Online repair needs to swap the contents of two files that are by definition
+   *not* identical.
+   For directory and xattr repairs, the user-visible contents might be the
+   same, but the contents of individual blocks may be very different.
+
+e. Old blocks in the file may be cross-linked with another structure and must
+   not reappear if the system goes down mid-repair.
+
+These problems are overcome by creating a new deferred operation and a new type
+of log intent item to track the progress of an operation to exchange two file
+ranges.
+The new deferred operation type chains together the same transactions used by
+the reverse-mapping extent swap code.
+The new log item records the progress of the exchange to ensure that once an
+exchange begins, it will always run to completion, even there are
+interruptions.
+
+The proposed patchset is the
+`atomic extent swap
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=atomic-file-updates>`_
+series.
+
+Using Log-Incompatible Feature Flags
+````````````````````````````````````
+
+Starting with XFS v5, the superblock contains a ``sb_features_log_incompat``
+field to indicate that the log contains records that might not readable by all
+kernels that could mount this filesystem.
+In short, log incompat features protect the log contents against kernels that
+will not understand the contents.
+Unlike the other superblock feature bits, log incompat bits are ephemeral
+because an empty (clean) log does not need protection.
+The log cleans itself after its contents have been committed into the
+filesystem, either as part of an unmount or because the system is otherwise
+idle.
+Because upper level code can be working on a transaction at the same time that
+the log cleans itself, it is necessary for upper level code to communicate to
+the log when it is going to use a log incompatible feature.
+
+The log coordinates access to incompatible features through the use of one
+``struct rw_semaphore`` for each feature.
+The log cleaning code tries to take this rwsem in exclusive mode to clear the
+bit; if the lock attempt fails, the feature bit remains set.
+Filesystem code signals its intention to use a log incompat feature in a
+transaction by calling ``xlog_use_incompat_feat``, which takes the rwsem in
+shared mode.
+The code supporting a log incompat feature should create wrapper functions to
+obtain the log feature and call ``xfs_add_incompat_log_feature`` to set the
+feature bits in the primary superblock.
+The superblock update is performed transactionally, so the wrapper to obtain
+log assistance must be called just prior to the creation of the transaction
+that uses the functionality.
+For a file operation, this step must happen after taking the IOLOCK and the
+MMAPLOCK, but before allocating the transaction.
+When the transaction is complete, the ``xlog_drop_incompat_feat`` function
+is called to release the feature.
+The feature bit will not be cleared from the superblock until the log becomes
+clean.
+
+Log-assisted extended attribute updates and atomic extent swaps both use log
+incompat features and provide convenience wrappers around the functionality.
+
+Mechanics of an Atomic Extent Swap
+``````````````````````````````````
+
+Swapping entire file forks is a complex task.
+The goal is to exchange all file fork mappings between two file fork offset
+ranges.
+There are likely to be many extent mappings in each fork, and the edges of
+the mappings aren't necessarily aligned.
+Furthermore, there may be other updates that need to happen after the swap,
+such as exchanging file sizes, inode flags, or conversion of fork data to local
+format.
+This is roughly the format of the new deferred extent swap work item:
+
+.. code-block:: c
+
+	struct xfs_swapext_intent {
+	    /* Inodes participating in the operation. */
+	    struct xfs_inode    *sxi_ip1;
+	    struct xfs_inode    *sxi_ip2;
+
+	    /* File offset range information. */
+	    xfs_fileoff_t       sxi_startoff1;
+	    xfs_fileoff_t       sxi_startoff2;
+	    xfs_filblks_t       sxi_blockcount;
+
+	    /* Set these file sizes after the operation, unless negative. */
+	    xfs_fsize_t         sxi_isize1;
+	    xfs_fsize_t         sxi_isize2;
+
+	    /* XFS_SWAP_EXT_* log operation flags */
+	    uint64_t            sxi_flags;
+	};
+
+The new log intent item contains enough information to track two logical fork
+offset ranges: ``(inode1, startoff1, blockcount)`` and ``(inode2, startoff2,
+blockcount)``.
+Each step of a swap operation exchanges the largest file range mapping possible
+from one file to the other.
+After each step in the swap operation, the two startoff fields are incremented
+and the blockcount field is decremented to reflect the progress made.
+The flags field captures behavioral parameters such as swapping the attr fork
+instead of the data fork and other work to be done after the extent swap.
+The two isize fields are used to swap the file size at the end of the operation
+if the file data fork is the target of the swap operation.
+
+When the extent swap is initiated, the sequence of operations is as follows:
+
+1. Create a deferred work item for the extent swap.
+   At the start, it should contain the entirety of the file ranges to be
+   swapped.
+
+2. Call ``xfs_defer_finish`` to start processing of the exchange.
+   This will log an extent swap intent item to the transaction for the deferred
+   extent swap work item.
+
+3. Until ``sxi_blockcount`` of the deferred extent swap work item is zero,
+
+   a. Read the block maps of both file ranges starting at ``sxi_startoff1`` and
+      ``sxi_startoff2``, respectively, and compute the longest extent that can
+      be swapped in a single step.
+      This is the minimum of the two ``br_blockcount`` s in the mappings.
+      Keep advancing through the file forks until at least one of the mappings
+      contains written blocks.
+      Mutual holes, unwritten extents, and extent mappings to the same physical
+      space are not exchanged.
+
+      For the next few steps, this document will refer to the mapping that came
+      from file 1 as "map1", and the mapping that came from file 2 as "map2".
+
+   b. Create a deferred block mapping update to unmap map1 from file 1.
+
+   c. Create a deferred block mapping update to unmap map2 from file 2.
+
+   d. Create a deferred block mapping update to map map1 into file 2.
+
+   e. Create a deferred block mapping update to map map2 into file 1.
+
+   f. Log the block, quota, and extent count updates for both files.
+
+   g. Extend the ondisk size of either file if necessary.
+
+   h. Log an extent swap done log item for the extent swap intent log item
+      that was read at the start of step 3.
+
+   i. Compute the amount of file range that has just been covered.
+      This quantity is ``(map1.br_startoff + map1.br_blockcount -
+      sxi_startoff1)``, because step 3a could have skipped holes.
+
+   j. Increase the starting offsets of ``sxi_startoff1`` and ``sxi_startoff2``
+      by the number of blocks computed in the previous step, and decrease
+      ``sxi_blockcount`` by the same quantity.
+      This advances the cursor.
+
+   k. Log a new extent swap intent log item reflecting the advanced state of
+      the work item.
+
+   l. Return the proper error code (EAGAIN) to the deferred operation manager
+      to inform it that there is more work to be done.
+      The operation manager completes the deferred work in steps 3b-3e before
+      moving back to the start of step 3.
+
+4. Perform any post-processing.
+   This will be discussed in more detail in subsequent sections.
+
+If the filesystem goes down in the middle of an operation, log recovery will
+find the most recent unfinished extent swap log intent item and restart from
+there.
+This is how extent swapping guarantees that an outside observer will either see
+the old broken structure or the new one, and never a mismash of both.
+
+Extent Swapping with Regular User Files
+```````````````````````````````````````
+
+As mentioned earlier, XFS has long had the ability to swap extents between
+files, which is used almost exclusively by ``xfs_fsr`` to defragment files.
+The earliest form of this was the fork swap mechanism, where the entire
+contents of data forks could be exchanged between two files by exchanging the
+raw bytes in each inode fork's immediate area.
+When XFS v5 came along with self-describing metadata, this old mechanism grew
+some log support to continue rewriting the owner fields of BMBT blocks during
+log recovery.
+When the reverse mapping btree was later added to XFS, the only way to maintain
+the consistency of the fork mappings with the reverse mapping index was to
+develop an iterative mechanism that used deferred bmap and rmap operations to
+swap mappings one at a time.
+This mechanism is identical to steps 2-3 from the procedure above except for
+the new tracking items, because the atomic extent swap mechanism is an
+iteration of an existing mechanism and not something totally novel.
+For the narrow case of file defragmentation, the file contents must be
+identical, so the recovery guarantees are not much of a gain.
+
+Atomic extent swapping is much more flexible than the existing swapext
+implementations because it can guarantee that the caller never sees a mix of
+old and new contents even after a crash, and it can operate on two arbitrary
+file fork ranges.
+The extra flexibility enables several new use cases:
+
+- **Atomic commit of file writes**: A userspace process opens a file that it
+  wants to update.
+  Next, it opens a temporary file and calls the file clone operation to reflink
+  the first file's contents into the temporary file.
+  Writes to the original file should instead be written to the temporary file.
+  Finally, the process calls the atomic extent swap system call
+  (``FIEXCHANGE_RANGE``) to exchange the file contents, thereby committing all
+  of the updates to the original file, or none of them.
+
+- **Transactional file updates**: The same mechanism as above, but the caller
+  only wants the commit to occur if the original file's contents have not
+  changed.
+  To make this happen, the calling process snapshots the file modification and
+  change timestamps of the original file before reflinking its data to the
+  temporary file.
+  When the program is ready to commit the changes, it passes the timestamps
+  into the kernel as arguments to the atomic extent swap system call.
+  The kernel only commits the changes if the provided timestamps match the
+  original file.
+
+- **Emulation of atomic block device writes**: Export a block device with a
+  logical sector size matching the filesystem block size to force all writes
+  to be aligned to the filesystem block size.
+  Stage all writes to a temporary file, and when that is complete, call the
+  atomic extent swap system call with a flag to indicate that holes in the
+  temporary file should be ignored.
+  This emulates an atomic device write in software, and can support arbitrary
+  scattered writes.
+
+Preparation for Extent Swapping
+```````````````````````````````
+
+There are a few things that need to be taken care of before initiating an
+atomic extent swap operation.
+First, regular files require the page cache to be flushed to disk before the
+operation begins, and directio writes to be quiesced.
+Like any filesystem operation, extent swapping must determine the maximum
+amount of disk space and quota that can be consumed on behalf of both files in
+the operation, and reserve that quantity of resources to avoid an unrecoverable
+out of space failure once it starts dirtying metadata.
+The preparation step scans the ranges of both files to estimate:
+
+- Data device blocks needed to handle the repeated updates to the fork
+  mappings.
+- Change in data and realtime block counts for both files.
+- Increase in quota usage for both files, if the two files do not share the
+  same set of quota ids.
+- The number of extent mappings that will be added to each file.
+- Whether or not there are partially written realtime extents.
+  User programs must never be able to access a realtime file extent that maps
+  to different extents on the realtime volume, which could happen if the
+  operation fails to run to completion.
+
+The need for precise estimation increases the run time of the swap operation,
+but it is very important to maintain correct accounting.
+The filesystem must not run completely out of free space, nor can the extent
+swap ever add more extent mappings to a fork than it can support.
+Regular users are required to abide the quota limits, though metadata repairs
+may exceed quota to resolve inconsistent metadata elsewhere.
+
+Special Features for Swapping Metadata File Extents
+```````````````````````````````````````````````````
+
+Extended attributes, symbolic links, and directories can set the fork format to
+"local" and treat the fork as a literal area for data storage.
+Metadata repairs must take extra steps to support these cases:
+
+- If both forks are in local format and the fork areas are large enough, the
+  swap is performed by copying the incore fork contents, logging both forks,
+  and committing.
+  The atomic extent swap mechanism is not necessary, since this can be done
+  with a single transaction.
+
+- If both forks map blocks, then the regular atomic extent swap is used.
+
+- Otherwise, only one fork is in local format.
+  The contents of the local format fork are converted to a block to perform the
+  swap.
+  The conversion to block format must be done in the same transaction that
+  logs the initial extent swap intent log item.
+  The regular atomic extent swap is used to exchange the mappings.
+  Special flags are set on the swap operation so that the transaction can be
+  rolled one more time to convert the second file's fork back to local format
+  if possible.
+
+Extended attributes and directories stamp the owning inode into every block,
+but the buffer verifiers do not actually check the inode number!
+Although there is no verification, it is still important to maintain
+referential integrity, so prior to performing the extent swap, online repair
+walks every block in the new data structure to update the owner field and flush
+the buffer to disk.
+
+After a successful swap operation, the repair operation must reap the old fork
+blocks by processing each fork mapping through the standard :ref:`file extent
+reaping <reaping>` mechanism that is done post-repair.
+If the filesystem should go down during the reap part of the repair, the
+iunlink processing at the end of recovery will free both the temporary file and
+whatever blocks were not reaped.
+However, this iunlink processing omits the cross-link detection of online
+repair, and is not completely foolproof.
+
+Swapping Temporary File Extents
+```````````````````````````````
+
+To repair a metadata file, online repair proceeds as follows:
+
+1. Create a temporary repair file.
+
+2. Use the staging data to write out new contents into the temporary repair
+   file.
+   The same fork must be written to as is being repaired.
+
+3. Commit the scrub transaction, since the swap estimation step must be
+   completed before transaction reservations are made.
+
+4. Call ``xrep_tempswap_trans_alloc`` to allocate a new scrub transaction with
+   the appropriate resource reservations, locks, and fill out a ``struct
+   xfs_swapext_req`` with the details of the swap operation.
+
+5. Call ``xrep_tempswap_contents`` to swap the contents.
+
+6. Commit the transaction to complete the repair.
+
+.. _rtsummary:
+
+Case Study: Repairing the Realtime Summary File
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In the "realtime" section of an XFS filesystem, free space is tracked via a
+bitmap, similar to Unix FFS.
+Each bit in the bitmap represents one realtime extent, which is a multiple of
+the filesystem block size between 4KiB and 1GiB in size.
+The realtime summary file indexes the number of free extents of a given size to
+the offset of the block within the realtime free space bitmap where those free
+extents begin.
+In other words, the summary file helps the allocator find free extents by
+length, similar to what the free space by count (cntbt) btree does for the data
+section.
+
+The summary file itself is a flat file (with no block headers or checksums!)
+partitioned into ``log2(total rt extents)`` sections containing enough 32-bit
+counters to match the number of blocks in the rt bitmap.
+Each counter records the number of free extents that start in that bitmap block
+and can satisfy a power-of-two allocation request.
+
+To check the summary file against the bitmap:
+
+1. Take the ILOCK of both the realtime bitmap and summary files.
+
+2. For each free space extent recorded in the bitmap:
+
+   a. Compute the position in the summary file that contains a counter that
+      represents this free extent.
+
+   b. Read the counter from the xfile.
+
+   c. Increment it, and write it back to the xfile.
+
+3. Compare the contents of the xfile against the ondisk file.
+
+To repair the summary file, write the xfile contents into the temporary file
+and use atomic extent swap to commit the new contents.
+The temporary file is then reaped.
+
+The proposed patchset is the
+`realtime summary repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rtsummary>`_
+series.
+
+Case Study: Salvaging Extended Attributes
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In XFS, extended attributes are implemented as a namespaced name-value store.
+Values are limited in size to 64KiB, but there is no limit in the number of
+names.
+The attribute fork is unpartitioned, which means that the root of the attribute
+structure is always in logical block zero, but attribute leaf blocks, dabtree
+index blocks, and remote value blocks are intermixed.
+Attribute leaf blocks contain variable-sized records that associate
+user-provided names with the user-provided values.
+Values larger than a block are allocated separate extents and written there.
+If the leaf information expands beyond a single block, a directory/attribute
+btree (``dabtree``) is created to map hashes of attribute names to entries
+for fast lookup.
+
+Salvaging extended attributes is done as follows:
+
+1. Walk the attr fork mappings of the file being repaired to find the attribute
+   leaf blocks.
+   When one is found,
+
+   a. Walk the attr leaf block to find candidate keys.
+      When one is found,
+
+      1. Check the name for problems, and ignore the name if there are.
+
+      2. Retrieve the value.
+         If that succeeds, add the name and value to the staging xfarray and
+         xfblob.
+
+2. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more attr fork blocks to examine, unlock the file and
+   add the staged extended attributes to the temporary file.
+
+3. Use atomic extent swapping to exchange the new and old extended attribute
+   structures.
+   The old attribute blocks are now attached to the temporary file.
+
+4. Reap the temporary file.
+
+The proposed patchset is the
+`extended attribute repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-xattrs>`_
+series.

[PATCH 12/14] xfs: document directory tree repairs

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Directory tree repairs are the least complete part of online fsck, due
to the lack of directory parent pointers.  However, even without that
feature, we can still make some corrections to the directory tree -- we
can salvage as many directory entries as we can from a damaged
directory, and we can reattach orphaned inodes to the lost+found, just
as xfs_repair does now.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  236 ++++++++++++++++++++
 1 file changed, 236 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index 6cdec62e3f23..cd26ccd3a0e4 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -4273,3 +4273,239 @@ The proposed patchset is the
 `extended attribute repair
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-xattrs>`_
 series.
+
+Fixing Directories
+------------------
+
+Fixing directories is difficult with currently available filesystem features.
+The offline repair tool scans all inodes to find files with nonzero link count,
+and then it scans all directories to establish parentage of those linked files.
+Damaged files and directories are zapped, and files with no parent are
+moved to the ``/lost+found`` directory.
+It does not try to salvage anything.
+
+The best that online repair can do at this time is to read directory data
+blocks and salvage any dirents that look plausible, correct link counts, and
+move orphans back into the directory tree.
+The salvage process is discussed in the case study at the end of this section.
+The second component to fixing the directory tree online is the :ref:`file link
+count fsck <nlinks>`, since it can scan the entire filesystem to make sure that
+files can neither be deleted while there are still parents nor forgotten after
+all parents sever their links to the child.
+The third part is discussed at the :ref:`end of this section<orphanage>`.
+However, there may be a solution to these deficiencies soon!
+
+Parent Pointers
+```````````````
+
+The lack of secondary directory metadata hinders directory tree reconstruction
+in much the same way that the historic lack of reverse space mapping
+information once hindered reconstruction of filesystem space metadata.
+Specifically, the lack of redundant metadata makes it nearly impossible to
+construct a true replacement for a damaged directory; the best repair can do is
+to salvage the dirents and use the file link count repair function to move
+orphaned files to the lost and found.
+The proposed parent pointer feature, however, will make total directory
+reconstruction possible.
+
+Directory parent pointers were first proposed as an XFS feature more than a
+decade ago by SGI.
+In that implementation, each link from a parent directory to a child file was
+augmented by an extended attribute in the child that could be used to identify
+the parent directory.
+Unfortunately, this early implementation had several major shortcomings:
+
+1. The XFS codebase of the late 2000s did not have the infrastructure to
+   enforce strong referential integrity in the directory tree, which is a fancy
+   way to say that it could not guarantee that a change in a forward link would
+   always be followed up by a corresponding change to the reverse links.
+
+2. Referential integrity was not integrated into either offline repair tool.
+   Checking had to be done online without taking any kernel or inode locks to
+   coordinate access.
+   It is not clear if this actually worked properly.
+
+3. The extended attribute did not record the name of the directory entry in the
+   parent, so the first parent pointer implementation cannot be used to
+   reconnect the directory tree.
+
+4. Extended attribute forks only support 65,536 extents, which means that
+   parent pointer attribute creation is likely to fail at some point before the
+   maximum file link count is achieved.
+
+In the second implementation (currently being developed by Allison Henderson
+and Chandan Babu), the extended attribute code will be enhanced to use log
+intent items to guarantee that an extended attribute update can always be
+completed by log recovery.
+The maximum extent counts of both the data and attribute forks have raised to
+allow for creation of as many parent pointers as possible.
+The parent pointer data will also include the entry name and location within
+the parent.
+In other words, child files will store parent pointer mappings of the form
+``(parent_ino, parent_gen, dirent_pos) → (dirent_name)`` in their extended
+attribute data.
+With that in place, XFS can guarantee strong referential integrity of directory
+tree operations -- forward links will always be complemented with reverse
+links.
+
+When the parent pointer feature lands, the directory checking process can be
+strengthened to ensure that the target of each dirent also contains a parent
+pointer pointing back to the dirent.
+The quality of directory repairs will improve because online fsck will be able
+to reconstruct a directory in its entirety instead of skipping unsalvageable
+areas.
+This process is imagined to involve a :ref:`coordinated inode scan <iscan>` and
+a :ref:`directory entry live update hook <liveupdate>`:
+Scan every file in the entire filesystem, and every time the scan encounters a
+file with a parent pointer to the directory that is being reconstructed, record
+this entry in the temporary directory.
+When the scan is complete, atomically swap the contents of the temporary
+directory and the directory being repaired.
+This code has not yet been constructed, so there is not yet a case study laying
+out exactly how this process works.
+
+Parent pointers themselves can be checked by scanning each pointer and
+verifying that the target of the pointer is a directory and that it contains a
+dirent that corresponds to the information recorded in the parent pointer.
+Reconstruction of the parent pointer information will work similarly to
+directory reconstruction -- scan the filesystem, record the dirents pointing to
+the file being repaired, and rebuild that part of the xattr namespace.
+
+**Question**: How will repair ensure that the ``dirent_pos`` fields match in
+the reconstructed directory?
+
+*Answer*: The field could be designated advisory, since the other three values
+are sufficient to find the entry in the parent.
+However, this makes indexed key lookup impossible while repairs are ongoing.
+A second option would be to allow creating directory entries at specified
+offsets, which solves the referential integrity problem but runs the risk that
+dirent creation will fail due to conflicts with the free space in the
+directory.
+These conflicts could be resolved by appending the directory entry and amending
+the xattr code to support updating an xattr key and reindexing the dabtree,
+though this would have to be performed with the parent directory still locked.
+A fourth option would be to remove the parent pointer entry and re-add it
+atomically.
+
+Case Study: Salvaging Directories
+`````````````````````````````````
+
+Unlike extended attributes, directory blocks are all the same size, so
+salvaging directories is straightforward:
+
+1. Find the parent of the directory.
+   If the dotdot entry is not unreadable, try to confirm that the alleged
+   parent has a child entry pointing back to the directory being repaired.
+   Otherwise, walk the filesystem to find it.
+
+2. Walk the first partition of data fork of the directory to find the directory
+   entry data blocks.
+   When one is found,
+
+   a. Walk the directory data block to find candidate entries.
+      When an entry is found:
+
+      i. Check the name for problems, and ignore the name if there are.
+
+      ii. Retrieve the inumber and grab the inode.
+          If that succeeds, add the name, inode number, and file type to the
+          staging xfarray and xblob.
+
+3. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more directory data blocks to examine, unlock the
+   directory and add the staged dirents into the temporary directory.
+   Truncate the staging files.
+
+4. Use atomic extent swapping to exchange the new and old directory structures.
+   The old directory blocks are now attached to the temporary file.
+
+5. Reap the temporary file.
+
+**Question**: Should repair invalidate dentries when rebuilding a directory?
+
+**Question**: Can the dentry cache know about a directory entry that cannot be
+salvaged?
+
+In theory, the dentry cache should be a subset of the directory entries on disk
+because there's no way to load a dentry without having something to read in the
+directory.
+However, it is possible for a coherency problem to be introduced if the ondisk
+structures becomes corrupt *after* the cache loads.
+In theory it is necessary to scan all dentry cache entries for a directory to
+ensure that one of the following apply:
+
+1. The cached dentry reflects an ondisk dirent in the new directory.
+
+2. The cached dentry no longer has a corresponding ondisk dirent in the new
+   directory and the dentry can be purged from the cache.
+
+3. The cached dentry no longer has an ondisk dirent but the dentry cannot be
+   purged.
+   This is bad.
+
+Unfortunately, the dentry cache does not have a means to walk all the dentries
+with a particular directory as a parent.
+This makes detecting situations #2 and #3 impossible, and remains an
+interesting question for research.
+
+The proposed patchset is the
+`directory repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-dirs>`_
+series.
+
+.. _orphanage:
+
+The Orphanage
+-------------
+
+Filesystems present files as a directed, and hopefully acyclic, graph.
+In other words, a tree.
+The root of the filesystem is a directory, and each entry in a directory points
+downwards either to more subdirectories or to non-directory files.
+Unfortunately, a disruption in the directory graph pointers result in a
+disconnected graph, which makes files impossible to access via regular path
+resolution.
+The directory parent pointer online scrub code can detect a dotdot entry
+pointing to a parent directory that doesn't have a link back to the child
+directory, and the file link count checker can detect a file that isn't pointed
+to by any directory in the filesystem.
+If the file in question has a positive link count, the file in question is an
+orphan.
+
+When orphans are found, they should be reconnected to the directory tree.
+Offline fsck solves the problem by creating a directory ``/lost+found`` to
+serve as an orphanage, and linking orphan files into the orphanage by using the
+inumber as the name.
+Reparenting a file to the orphanage does not reset any of its permissions or
+ACLs.
+
+This process is more involved in the kernel than it is in userspace.
+The directory and file link count repair setup functions must use the regular
+VFS mechanisms to create the orphanage directory with all the necessary
+security attributes and dentry cache entries, just like a regular directory
+tree modification.
+
+Orphaned files are adopted by the orphanage as follows:
+
+1. Call ``xrep_orphanage_try_create`` at the start of the scrub setup function
+   to try to ensure that the lost and found directory actually exists.
+   This also attaches the orphanage directory to the scrub context.
+
+2. If the decision is made to reconnect a file, take the IOLOCK of both the
+   orphanage and the file being reattached.
+   The ``xrep_orphanage_iolock_two`` function follows the inode locking
+   strategy discussed earlier.
+
+3. Call ``xrep_orphanage_compute_blkres`` and ``xrep_orphanage_compute_name``
+   to compute the new name in the orphanage and the block reservation required.
+
+4. Use ``xrep_orphanage_adoption_prep`` to reserve resources to the repair
+   transaction.
+
+5. Call ``xrep_orphanage_adopt`` to reparent the orphaned file into the lost
+   and found, and update the kernel dentry cache.
+
+The proposed patches are in the
+`orphanage adoption
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-orphanage>`_
+series.

[PATCH 13/14] xfs: document the userspace fsck driver program

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Add the sixth chapter of the online fsck design documentation, where
we discuss the details of the data structures and algorithms used by the
driver program xfs_scrub.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  313 ++++++++++++++++++++
 1 file changed, 313 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index cd26ccd3a0e4..adffc8c31c9c 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -300,6 +300,9 @@ The seven phases are as follows:
 7. Re-check the summary counters and presents the caller with a summary of
    space usage and file counts.
 
+This allocation of responsibilities will be :ref:`revisited <scrubcheck>`
+later in this document.
+
 Steps for Each Scrub Item
 -------------------------
 
@@ -4509,3 +4512,313 @@ The proposed patches are in the
 `orphanage adoption
 <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-orphanage>`_
 series.
+
+6. Userspace Algorithms and Data Structures
+===========================================
+
+This section discusses the key algorithms and data structures of the userspace
+program, ``xfs_scrub``, that provide the ability to drive metadata checks and
+repairs in the kernel, verify file data, and look for other potential problems.
+
+.. _scrubcheck:
+
+Checking Metadata
+-----------------
+
+Recall the :ref:`phases of fsck work<scrubphases>` outlined earlier.
+That structure follows naturally from the data dependencies designed into the
+filesystem from its beginnings in 1993.
+In XFS, there are several groups of metadata dependencies:
+
+a. Filesystem summary counts depend on consistency within the inode indices,
+   the allocation group space btrees, and the realtime volume space
+   information.
+
+b. Quota resource counts depend on consistency within the quota file data
+   forks, inode indices, inode records, and the forks of every file on the
+   system.
+
+c. The naming hierarchy depends on consistency within the directory and
+   extended attribute structures.
+   This includes file link counts.
+
+d. Directories, extended attributes, and file data depend on consistency within
+   the file forks that map directory and extended attribute data to physical
+   storage media.
+
+e. The file forks depends on consistency within inode records and the space
+   metadata indices of the allocation groups and the realtime volume.
+   This includes quota and realtime metadata files.
+
+f. Inode records depends on consistency within the inode metadata indices.
+
+g. Realtime space metadata depend on the inode records and data forks of the
+   realtime metadata inodes.
+
+h. The allocation group metadata indices (free space, inodes, reference count,
+   and reverse mapping btrees) depend on consistency within the AG headers and
+   between all the AG metadata btrees.
+
+i. ``xfs_scrub`` depends on the filesystem being mounted and kernel support
+   for online fsck functionality.
+
+Therefore, a metadata dependency graph is a convenient way to schedule checking
+operations in the ``xfs_scrub`` program:
+
+- Phase 1 checks that the provided path maps to an XFS filesystem and detect
+  the kernel's scrubbing abilities, which validates group (i).
+
+- Phase 2 scrubs groups (g) and (h) in parallel using a threaded workqueue.
+
+- Phase 3 checks groups (f), (e), and (d), in that order.
+  These groups are all file metadata, which means that inodes are scanned in
+  parallel.
+
+- Phase 4 repairs everything in groups (i) through (d) so that phases 5 and 6
+  may run reliably.
+
+- Phase 5 starts by checking groups (b) and (c) in parallel before moving on
+  to checking names.
+
+- Phase 6 depends on groups (i) through (b) to find file data blocks to verify,
+  to read them, and to report which blocks of which files are affected.
+
+- Phase 7 checks group (a), having validated everything else.
+
+Notice that the data dependencies between groups are enforced by the structure
+of the program flow.
+
+Parallel Inode Scans
+--------------------
+
+An XFS filesystem can easily contain hundreds of millions of inodes.
+Given that XFS targets installations with large high-performance storage,
+it is desirable to scrub inodes in parallel to minimize runtime, particularly
+if the program has been invoked manually from a command line.
+This requires careful scheduling to keep the threads as evenly loaded as
+possible.
+
+Early iterations of the ``xfs_scrub`` inode scanner naïvely created a single
+workqueue and scheduled a single workqueue item per AG.
+Each workqueue item walked the inode btree (with ``XFS_IOC_INUMBERS``) to find
+inode chunks and then called bulkstat (``XFS_IOC_BULKSTAT``) to gather enough
+information to construct file handles.
+The file handle was then passed to a function to generate scrub items for each
+metadata object of each inode.
+This simple algorithm leads to thread balancing problems in phase 3 if the
+filesystem contains one AG with a few large sparse files and the rest of the
+AGs contain many smaller files.
+The inode scan dispatch function was not sufficiently granular; it should have
+been dispatching at the level of individual inodes, or, to constrain memory
+consumption, inode btree records.
+
+Thanks to Dave Chinner, bounded workqueues in userspace enable ``xfs_scrub`` to
+avoid this problem with ease by adding a second workqueue.
+Just like before, the first workqueue is seeded with one workqueue item per AG,
+and it uses INUMBERS to find inode btree chunks.
+The second workqueue, however, is configured with an upper bound on the number
+of items that can be waiting to be run.
+Each inode btree chunk found by the first workqueue's workers are queued to the
+second workqueue, and it is this second workqueue that queries BULKSTAT,
+creates a file handle, and passes it to a function to generate scrub items for
+each metadata object of each inode.
+If the second workqueue is too full, the workqueue add function blocks the
+first workqueue's workers until the backlog eases.
+This doesn't completely solve the balancing problem, but reduces it enough to
+move on to more pressing issues.
+
+The proposed patchsets are the scrub
+`performance tweaks
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-performance-tweaks>`_
+and the
+`inode scan rebalance
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-iscan-rebalance>`_
+series.
+
+.. _scrubrepair:
+
+Scheduling Repairs
+------------------
+
+During phase 2, corruptions and inconsistencies reported in any AGI header or
+inode btree are repaired immediately, because phase 3 relies on proper
+functioning of the inode indices to find inodes to scan.
+Failed repairs are rescheduled to phase 4.
+Problems reported in any other space metadata are deferred to phase 4.
+Optimization opportunities are always deferred to phase 4, no matter their
+origin.
+
+During phase 3, corruptions and inconsistencies reported in any part of a
+file's metadata are repaired immediately if all space metadata were validated
+during phase 2.
+Repairs that fail or cannot be repaired immediately are scheduled for phase 4.
+
+In the original design of ``xfs_scrub``, it was thought that repairs would be
+so infrequent that the ``struct xfs_scrub_metadata`` objects used to
+communicate with the kernel could also be used as the primary object to
+schedule repairs.
+With recent increases in the number of optimizations possible for a given
+filesystem object, it became much more memory-efficient to track all eligible
+repairs for a given filesystem object with a single repair item.
+Each repair item represents a single lockable object -- AGs, metadata files,
+individual inodes, or a class of summary information.
+
+Phase 4 is responsible for scheduling a lot of repair work in as quick a
+manner as is practical.
+The :ref:`data dependencies <scrubcheck>` outlined earlier still apply, which
+means that ``xfs_scrub`` must try to complete the repair work scheduled by
+phase 2 before trying repair work scheduled by phase 3.
+The repair process is as follows:
+
+1. Start a round of repair with a workqueue and enough workers to keep the CPUs
+   as busy as the user desires.
+
+   a. For each repair item queued by phase 2,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   b. If any repairs were made, jump back to 1a to retry all the phase 2 items.
+
+   c. For each repair item queued by phase 3,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   d. If any repairs were made, jump back to 1c to retry all the phase 3 items.
+
+2. If step 1 made any repair progress of any kind, jump back to step 1 to start
+   another round of repair.
+
+3. If there are items left to repair, run them all serially one more time.
+   Complain if the repairs were not successful, since this is the last chance
+   to repair anything.
+
+Corruptions and inconsistencies encountered during phases 5 and 7 are repaired
+immediately.
+Corrupt file data blocks reported by phase 6 cannot be recovered by the
+filesystem.
+
+The proposed patchsets are the
+`repair warning improvements
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-better-repair-warnings>`_,
+refactoring of the
+`repair data dependency
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-repair-data-deps>`_
+and
+`object tracking
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-object-tracking>`_,
+and the
+`repair scheduling
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-repair-scheduling>`_
+improvement series.
+
+Checking Names for Confusable Unicode Sequences
+-----------------------------------------------
+
+If ``xfs_scrub`` succeeds in validating the filesystem metadata by the end of
+phase 4, it moves on to phase 5, which checks for suspicious looking names in
+the filesystem.
+These names consist of the filesystem label, names in directory entries, and
+the names of extended attributes.
+Like most Unix filesystems, XFS imposes the sparest of constraints on the
+contents of a name -- slashes and null bytes are not allowed in directory
+entries; and null bytes are not allowed in extended attributes and the
+filesystem label.
+Directory entries and attribute keys store the length of the name explicitly
+ondisk, which means that nulls are not name terminators.
+For this section, the term "naming domain" refers to any place where names are
+presented together -- all the names in a directory, or all the attributes of a
+file.
+
+Although the Unix naming constraints are very permissive, the reality of most
+modern-day Linux systems is that programs work with Unicode character code
+points to support international languages.
+These programs typically encode those code points in UTF-8 when interfacing
+with the C library because the kernel expects null-terminated names.
+In the common case, therefore, names found in an XFS filesystem are actually
+UTF-8 encoded Unicode data.
+
+To maximize its expressiveness, the Unicode standard defines separate control
+points for various characters that render similarly or identically in writing
+systems around the world.
+For example, the character "Cyrillic Small Letter A" U+0430 "а" often renders
+identically to "Latin Small Letter A" U+0061 "a".
+
+The standard also permits characters to be constructed in multiple ways --
+either by using a defined code point, or by combining one code point with
+various combining marks.
+For example, the character "Angstrom Sign U+212B "Å" can also be expressed
+as "Latin Capital Letter A" U+0041 "A" followed by "Combining Ring Above"
+U+030A "◌̊".
+Both sequences render identically.
+
+Like the standards that preceded it, Unicode also defines various control
+characters to alter the presentation of text.
+For example, the character "Right-to-Left Override" U+202E can trick some
+programs into rendering "moo\\xe2\\x80\\xaegnp.txt" as "mootxt.png".
+A second category of rendering problems involves whitespace characters.
+If the character "Zero Width Space" U+200B is encountered in a file name, the
+name will render identically to a name that does not have the zero width
+space.
+
+If two names within a naming domain have different byte sequences but render
+identically, a user may be confused by it.
+The kernel, in its indifference to upper level encoding schemes, permits this.
+Most filesystem drivers persist the byte sequence names that are given to them
+by the VFS.
+
+Techniques for detecting confusable names are explained in great detail in
+sections 4 and 5 of the
+`Unicode Security Mechanisms <https://unicode.org/reports/tr39/>`_
+document.
+``xfs_scrub``, when it detects UTF-8 encoding in use on a system, uses the
+Unicode normalization form NFD in conjunction with the confusable name
+detection component of
+`libicu <https://github.com/unicode-org/icu>`_
+to identify names with a directory or within a file's extended attributes that
+could be confused for each other.
+Names are also checked for control characters, non-rendering characters, and
+mixing of bidirectional characters.
+All of these potential issues are reported to the system administrator during
+phase 5.
+
+Media Verification of File Data Extents
+---------------------------------------
+
+The system administrator can elect to initiate a media scan of all file data
+blocks.
+This scan after validation of all filesystem metadata (except for the summary
+counters) as phase 6.
+The scan starts by calling ``FS_IOC_GETFSMAP`` to scan the filesystem space map
+to find areas that are allocated to file data fork extents.
+Gaps betweeen data fork extents that are smaller than 64k are treated as if
+they were data fork extents to reduce the command setup overhead.
+When the space map scan accumulates a region larger than 32MB, a media
+verification request is sent to the disk as a directio read of the raw block
+device.
+
+If the verification read fails, ``xfs_scrub`` retries with single-block reads
+to narrow down the failure to the specific region of the media and recorded.
+When it has finished issuing verification requests, it again uses the space
+mapping ioctl to map the recorded media errors back to metadata structures
+and report what has been lost.
+For media errors in blocks owned by files, the lack of parent pointers means
+that the entire filesystem must be walked to report the file paths and offsets
+corresponding to the media error.

[PATCH 14/14] xfs: document future directions of online fsck

[Darrick J. Wong]

From: Darrick J. Wong <djwong@kernel.org>

Add the seventh and final chapter of the online fsck documentation,
where we talk about future functionality that can tie in with the
functionality provided by the online fsck patchset.

Signed-off-by: Darrick J. Wong <djwong@kernel.org>
---
 .../filesystems/xfs-online-fsck-design.rst         |  155 ++++++++++++++++++++
 1 file changed, 155 insertions(+)


diff --git a/Documentation/filesystems/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs-online-fsck-design.rst
index adffc8c31c9c..3faaae0038a6 100644
--- a/Documentation/filesystems/xfs-online-fsck-design.rst
+++ b/Documentation/filesystems/xfs-online-fsck-design.rst
@@ -4071,6 +4071,8 @@ The extra flexibility enables several new use cases:
   (``FIEXCHANGE_RANGE``) to exchange the file contents, thereby committing all
   of the updates to the original file, or none of them.
 
+.. _swapext_if_unchanged:
+
 - **Transactional file updates**: The same mechanism as above, but the caller
   only wants the commit to occur if the original file's contents have not
   changed.
@@ -4822,3 +4824,156 @@ and report what has been lost.
 For media errors in blocks owned by files, the lack of parent pointers means
 that the entire filesystem must be walked to report the file paths and offsets
 corresponding to the media error.
+
+7. Conclusion and Future Work
+=============================
+
+It is hoped that the reader of this document has followed the designs laid out
+in this document and now has some familiarity with how XFS performs online
+rebuilding of its metadata indices, and how filesystem users can interact with
+that functionality.
+Although the scope of this work is daunting, it is hoped that this guide will
+make it easier for code readers to understand what has been built, for whom it
+has been built, and why.
+Please feel free to contact the XFS mailing list with questions.
+
+FIEXCHANGE_RANGE
+----------------
+
+As discussed earlier, a second frontend to the atomic extent swap mechanism is
+a new ioctl call that userspace programs can use to commit updates to files
+atomically.
+This frontend has been out for review for several years now, though the
+necessary refinements to online repair and lack of customer demand mean that
+the proposal has not been pushed very hard.
+
+Vectorized Scrub
+----------------
+
+As it turns out, the :ref:`refactoring <scrubrepair>` of repair items mentioned
+earlier was a catalyst for enabling a vectorized scrub system call.
+Since 2018, the cost of making a kernel call has increased considerably on some
+systems to mitigate the effects of speculative execution attacks.
+This incentivizes program authors to make as few system calls as possible to
+reduce the number of times an execution path crosses a security boundary.
+
+With vectorized scrub, userspace pushes to the kernel the identity of a
+filesystem object, a list of scrub types to run against that object, and a
+simple representation of the data dependencies between the selected scrub
+types.
+The kernel executes as much of the caller's plan as it can until it hits a
+dependency that cannot be satisfied due to a corruption, and tells userspace
+how much was accomplished.
+It is hoped that ``io_uring`` will pick up enough of this functionality that
+online fsck can use that instead of adding a separate vectored scrub system
+call to XFS.
+
+The relevant patchsets are the
+`kernel vectorized scrub
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=vectorized-scrub>`_
+and
+`userspace vectorized scrub
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=vectorized-scrub>`_
+series.
+
+Quality of Service Targets for Scrub
+------------------------------------
+
+One serious shortcoming of the online fsck code is that the amount of time that
+it can spend in the kernel holding resource locks is basically unbounded.
+Userspace is allowed to send a fatal signal to the process which will cause
+``xfs_scrub`` to exit when it reaches a good stopping point, but there's no way
+for userspace to provide a time budget to the kernel.
+Given that the scrub codebase has helpers to detect fatal signals, it shouldn't
+be too much work to allow userspace to specify a timeout for a scrub/repair
+operation and abort the operation if it exceeds budget.
+However, most repair functions have the property that once they begin to touch
+ondisk metadata, the operation cannot be cancelled cleanly, after which a QoS
+timeout is no longer useful.
+
+Defragmenting Free Space
+------------------------
+
+Over the years, many XFS users have requested the creation of a program to
+clear a portion of the physical storage underlying a filesystem so that it
+becomes a contiguous chunk of free space.
+Call this free space defragmenter ``clearspace`` for short.
+
+The first piece the ``clearspace`` program needs is the ability to read the
+reverse mapping index from userspace.
+This already exists in the form of the ``FS_IOC_GETFSMAP`` ioctl.
+The second piece it needs is a new fallocate mode
+(``FALLOC_FL_MAP_FREE_SPACE``) that allocates the free space in a region and
+maps it to a file.
+Call this file the "space collector" file.
+The third piece is the ability to force an online repair.
+
+To clear all the metadata out of a portion of physical storage, clearspace
+uses the new fallocate map-freespace call to map any free space in that region
+to the space collector file.
+Next, clearspace finds all metadata blocks in that region by way of
+``GETFSMAP`` and issues forced repair requests on the data structure.
+This often results in the metadata being rebuilt somewhere that is not being
+cleared.
+After each relocation, clearspace calls the "map free space" function again to
+collect any newly freed space in the region being cleared.
+
+To clear all the file data out of a portion of the physical storage, clearspace
+uses the FSMAP information to find relevant file data blocks.
+Having identified a good target, it uses the ``FICLONERANGE`` call on that part
+of the file to try to share the physical space with a dummy file.
+Cloning the extent means that the original owners cannot overwrite the
+contents; any changes will be written somewhere else via copy-on-write.
+Clearspace makes its own copy of the frozen extent in an area that is not being
+cleared, and uses ``FIEDEUPRANGE`` (or the :ref:`atomic extent swap
+<swapext_if_unchanged>` feature) to change the target file's data extent
+mapping away from the area being cleared.
+When all other mappings have been moved, clearspace reflinks the space into the
+space collector file so that it becomes unavailable.
+
+There are further optimizations that could apply to the above algorithm.
+To clear a piece of physical storage that has a high sharing factor, it is
+strongly desirable to retain this sharing factor.
+In fact, these extents should be moved first to maximize sharing factor after
+the operation completes.
+To make this work smoothly, clearspace needs a new ioctl
+(``FS_IOC_GETREFCOUNTS``) to report reference count information to userspace.
+With the refcount information exposed, clearspace can quickly find the longest,
+most shared data extents in the filesystem, and target them first.
+
+**Question**: How might the filesystem move inode chunks?
+
+*Answer*: Dave Chinner has a prototype that creates a new file with the old
+contents and then locklessly runs around the filesystem updating directory
+entries.
+The operation cannot complete if the filesystem goes down.
+That problem isn't totally insurmountable: create an inode remapping table
+hidden behind a jump label, and a log item that tracks the kernel walking the
+filesystem to update directory entries.
+The trouble is, the kernel can't do anything about open files, since it cannot
+revoke them.
+
+**Question**: Can static keys be used to add a revoke bailout return to
+*every* code path coming in from userspace?
+
+*Answer*: In principle, yes.
+This would eliminate the overhead of the check until a revocation happens.
+It's not clear what we do to a revoked file after all the callers are finished
+with it, however.
+
+The relevant patchsets are the
+`kernel freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=defrag-freespace>`_
+and
+`userspace freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=defrag-freespace>`_
+series.
+
+Shrinking Filesystems
+---------------------
+
+Removing the end of the filesystem ought to be a simple matter of evacuating
+the data and metadata at the end of the filesystem, and handing the freed space
+to the shrink code.
+That requires an evacuation of the space at end of the filesystem, which is a
+use of free space defragmentation!