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 decision making.

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.

Table of Contents

1. What is a Filesystem Check?

A Unix filesystem has four main responsibilities:

  • Provide a hierarchy of names through which application programs can associate arbitrary blobs of data for any length of time,

  • Virtualize physical storage media across those names, and

  • Retrieve the named data blobs at any time.

  • Examine resource usage.

Metadata directly supporting these functions (e.g. files, directories, space mappings) are sometimes called primary metadata. Secondary metadata (e.g. reverse mapping and directory parent pointers) support operations internal to the filesystem, such as internal consistency checking and reorganization. Summary metadata, as the name implies, condense information contained in primary metadata for performance reasons.

The filesystem check (fsck) tool examines all the metadata in a filesystem to look for errors. In addition to looking for obvious metadata corruptions, fsck also cross-references different types of metadata records with each other to look for inconsistencies. People do not like losing data, so most fsck tools also contains some ability to correct 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. 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.


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 arrays. fsck tools address only the first problem.

TLDR; Show Me the Code!

Code is posted to the kernel.org git trees as follows: kernel changes, userspace changes, and QA test changes. 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 the filesystem when unexpected shutdowns occur as a result of silent corruptions in the 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 performed by the storage system administrator 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.


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".

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. Sharding enables better performance on highly parallel systems and helps to contain the damage when corruptions occur. 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.

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.


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.


In principle, online fsck should be able to check and to repair everything that the offline fsck program can handle. However, online fsck cannot be running 100% of the time, which means that latent errors may creep in after a scrub completes. If these errors cause the next mount to fail, offline fsck is the only solution. 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 is not a complete replacement for offline fsck, and a complete run of online fsck may take longer than online fsck. 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.

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 using the information in the scrub item to resubmit the kernel scrub call 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 and xfs_scrub is permitted to perform repairs, and there were no problems detected during phase 2, then those scrub items are repaired immediately. Optimizations, deferred repairs, 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.

This allocation of responsibilities will be revisited later in this document.

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:

  • Free space and reference count information

  • Inode records and indexes

  • Storage mapping information for file data

  • Directories

  • Extended attributes

  • Symbolic links

  • Quota limits

Scrub obeys the same rules as regular filesystem accesses for resource and lock acquisition.

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", 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.

Secondary metadata include:

  • Reverse mapping information

  • Directory parent pointers

This class of metadata is 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. Frequently, this means that full filesystems scans are necessary to rebuild the metadata. 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 will either have the same format as the ondisk structure or 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. While the filesystem scan is in progress, the repair function hooks the filesystem so that it can apply pending filesystem updates to the staging information. 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", 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.

Future Work Question: Can the full scan and live update code used to facilitate a repair also be used to implement a comprehensive check?

Answer: In theory, yes. Check would be much stronger if each scrub function employed these live scans to build a shadow copy of the metadata and then compared the shadow records to the ondisk records. However, doing that is a fair amount more work than what the checking functions do now. The live scans and hooks were developed much later. That in turn increases the runtime of those scrub functions.

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.

Examples of summary information include:

  • Summary counts of free space and inodes

  • File link counts from directories

  • Quota resource usage counts

Check and repair 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", 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.

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, 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

The testing plan for online fsck includes extending the existing fs testing infrastructure 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, fuzzing baselines, and improvements in fuzz testing comprehensiveness.

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 and the evolution of existing per-function stress testing.

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 timers and services that run online fsck automatically on weekends by default. 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. This can be tuned by the systemd administrator at any time to suit the latency and throughput requirements of customer 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.

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.

Future Work Question: Should the health reporting integrate with the new inotify fs error notification system? 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 and preservation of sickness info during memory reclaim.

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.

Whenever a file system operation modifies a block, the change is submitted to the log as part of a transaction. The log then processes these transactions marking them done once they are safely persisted to storage. 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 discrepancies that can be introduced 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 XFS Self Describing Metadata

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 encoding techniques.

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. Like any system that improves redundancy, the reverse-mapping feature increases overhead costs for space mapping activities. However, it has two critical advantages: first, the reverse index is key to enabling online fsck and other requested functionality such as free space defragmentation, better media failure reporting, and filesystem 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 by turning 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:

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.

There are several observations to make about reverse mapping indices:

  1. Reverse mappings can provide a positive affirmation of correctness if any of the above primary metadata are in doubt. The checking code for most primary metadata follows a path similar to the one outlined above.

  2. Proving the consistency of secondary metadata with the primary metadata is difficult because that requires a full scan of all primary space metadata, which is very time intensive. For example, checking a reverse mapping record for a file extent mapping btree block requires locking the file and searching the entire btree to confirm the block. Instead, scrub relies on rigorous cross-referencing during the primary space mapping structure checks.

  3. Consistency scans must use non-blocking lock acquisition primitives if the required locking order is not the same order used by regular filesystem operations. For example, if the filesystem normally takes a file ILOCK before taking the AGF buffer lock but scrub wants to take a file ILOCK while holding an AGF buffer lock, scrub cannot block on that second acquisition. This means that forward progress during this part of a scan of the reverse mapping data cannot be guaranteed if system load is heavy.

In summary, reverse mappings play a key role in reconstruction of primary metadata. 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

After the buffer cache, the next 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. 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). 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.

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

After internal block checks, 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, inode btree, and rmap btree records; to find mergeable records; and to improve cross referencing with rmap 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:

    1. Validate that the name does not contain invalid characters.

    2. 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:

    1. Does the name contain no invalid characters?

    2. Does the inumber correspond to an actual, allocated inode?

    3. Does the child inode have a nonzero link count?

    4. If a file type is included in the dirent, does it match the type of the inode?

    5. If the child is a subdirectory, does the child's dotdot pointer point back to the parent?

    6. 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 parents and file link counts 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 more detail later.

Quota usage and file link count 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.

Eventual Consistency vs. Online Fsck

Complex operations can make modifications to multiple per-AG data structures with a chain of transactions. These chains, once committed to the log, are restarted during log recovery if the system crashes while processing the chain. Because the AG header buffers are unlocked between transactions within a chain, online checking must coordinate with chained operations that are in progress to avoid incorrectly detecting inconsistencies due to pending chains. Furthermore, online repair must not run when operations are pending because the metadata are temporarily inconsistent with each other, and rebuilding is not possible.

Only online fsck has this requirement of total consistency of AG metadata, and should be relatively rare as compared to filesystem change operations. Online fsck coordinates with transaction chains as follows:

  • For each AG, maintain a count of intent items targetting that AG. The count should be bumped whenever a new item is added to the chain. The count should be dropped when the filesystem has locked the AG header buffers and finished the work.

  • 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. If the count is zero, proceed with the checking operation. If it is nonzero, cycle the buffer locks to allow the chain to make forward progress.

This may lead to online fsck taking a long time to complete, but regular filesystem updates take precedence over background checking activity. Details about the discovery of this situation are presented in the next section, and details about the solution are presented after that.

Discovery of the Problem

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 subtleties to XFS' transaction chaining strategy to consider:

  • 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.

  • AG header buffers are 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 other solutions to this problem were evaluated upon discovery of this flaw and rejected:

  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.

Intent Drains

Online fsck uses an atomic intent item counter and lock cycling to coordinate with transaction chains. 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. Call the appropriate subsystem function to add a deferred work item to a transaction.

  2. The function calls xfs_defer_drain_bump to increase the counter.

  3. When the deferred item manager wants to finish the deferred work item, it calls ->finish_item to complete it.

  4. The ->finish_item implementation logs some changes and calls xfs_defer_drain_drop to decrease the sloppy counter and wake up any threads waiting on the drain.

  5. 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 counter is zero (xfs_defer_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_defer_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 to zero.

The proposed patchset is the scrub intent drain series.

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_fsgates_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_fsgates_enable.

For more information, please see the kernel documentation of Static Keys.

Pageable Kernel Memory

Some online checking functions work by scanning the filesystem to build a shadow copy of an ondisk metadata structure in memory and comparing the two copies. For online repair to rebuild a metadata structure, it must compute the record set that will be stored in the new structure before it can persist that new structure to disk. Ideally, repairs complete with a single atomic commit that introduces a new data structure. To meet these goals, the kernel needs to collect a large amount of information in a place that doesn't require the correct operation of the filesystem.

Kernel memory isn't suitable because:

  • Allocating a contiguous region of memory to create a C array is very difficult, especially on 32-bit systems.

  • Linked lists of records introduce double pointer overhead which is very high and eliminate the possibility of indexed lookups.

  • Kernel memory is pinned, which can drive the system into OOM conditions.

  • The system might not have sufficient memory to stage all the information.

At any given time, online fsck does not need to keep the entire record set in memory, which means that individual records can be paged out if necessary. Continued development of online fsck demonstrated that the ability to perform indexed data storage would also be very useful. Fortunately, the Linux kernel already has a facility for byte-addressable and pageable storage: tmpfs. In-kernel graphics drivers (most notably i915) take advantage of tmpfs 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!

Historical Sidebar:

The first edition of online repair inserted records into a new btree as it found them, which failed because filesystem could shut down with a built data structure, which would be live after recovery finished.

The second edition solved the half-rebuilt structure problem by storing everything in memory, but frequently ran the system out of memory.

The third edition solved the OOM problem by using linked lists, but the memory overhead of the list pointers was extreme.

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 realtime summary 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 is convenient to access storage directly with pointers, just like userspace code does with regular memory. Online fsck must not drive the system into OOM conditions, which means that xfiles must be responsive to memory reclamation. tmpfs can only push a pagecache folio to the swap cache if the folio is neither pinned nor locked, which means the xfile must not pin too many folios.

Short term direct access to xfile contents is done by locking the pagecache folio and mapping it into kernel address space. Programmatic access (e.g. pread and pwrite) uses this mechanism. Folio locks are not supposed to be held for long periods of time, so long term direct access to xfile contents is done by bumping the folio refcount, mapping it into kernel address space, and dropping the folio lock. These long term users must be responsive to memory reclaim by hooking into the shrinker infrastructure to know when to release folios.

The xfile_get_page and xfile_put_page functions are provided to retrieve the (locked) folio that backs part of an xfile and to release it. The only code to use these folio lease functions are the xfarray sorting algorithms and the in-memory btrees.

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 writers 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 readers 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. tmpfs cannot be frozen, and xfiles must not be exposed to userspace.

If an xfile is shared between threads to stage repairs, the caller must provide its own locks to coordinate access. For example, if a scrub function stores scan results in an xfile and needs other threads to provide updates to the scanned data, the scrub function must provide a lock for all threads to share.

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.

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.

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:

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.

xfarray_idx_t i = XFARRAY_CURSOR_INIT;
while ((ret = xfarray_iter(array, &i, &rec)) == 1) {
    /* do something with rec */
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.

Case Study: Sorting xfarrays

The sorting algorithm used in the xfarray is actually a combination of adaptive quicksort and a heapsort subalgorithm in the spirit of Sedgewick and 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(n2) 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 series.

In-Memory B+Trees

The chapter about 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 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 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 next section 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.

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.

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:

default_load_factor = (maxrecs + minrecs) / 2

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 ? default_load_factor : 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 ? default_load_factor : 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.

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 reaping for more details.

Proposed patchsets are the bitmap rework and the preparation for bulk loading btrees.

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:

│RRR │

Sibling pointers are set every time a new block is added to the level:

┌────┐ ┌────┐ ┌────┐ ┌────┐
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘

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:

    ┌────┐       ┌────┐
    │PP  │←──────│PP  │
    └────┘       └────┘
    ↙   ↘         ↙   ↘
┌────┐ ┌────┐ ┌────┐ ┌────┐
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘

When it reaches the root level, it is ready to commit the new btree!:

        │  root   │
        │   PP    │
        ↙         ↘
    ┌────┐       ┌────┐
    │PP  │←──────│PP  │
    └────┘       └────┘
    ↙   ↘         ↙   ↘
┌────┐ ┌────┐ ┌────┐ ┌────┐
│RRR │←│RRR │←│RRR │←│RRR │
└────┘ └────┘ └────┘ └────┘

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:

    1. 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.

    2. 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.

    3. 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 section 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 series.

Case Study: Rebuilding the Space Reference Counts

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.

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 because these are extents allocated to stage a copy on write operation and are tracked in the refcount btree.

    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. This matches the sorting order of records in the refcount btree.

  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.

  • Until the reverse mapping btree runs out of records:

    • Retrieve the next record from the btree and put it in a bag.

    • Collect all records with the same starting block from the btree and put them in the bag.

    • While the bag isn't empty:

      • Among the mappings in the bag, compute the lowest block number where the reference count changes. This position will be either the starting block number of the next unprocessed reverse mapping or the next block after the shortest mapping in the bag.

      • Remove all mappings from the bag that end at this position.

      • Collect all reverse mappings that start at this position from the btree and put them in the bag.

      • If the size of the bag changed and is greater than one, create a new refcount record associating the block number range that we just walked to the size of the bag.

The bag-like structure in this case is a type 2 xfarray as discussed in the 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 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 mapping repair series.

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:

  1. 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.

  2. 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.

  3. If the region is crosslinked, delete the reverse mapping entry for the structure being repaired and move on to the next region.

  4. If the region is to be freed, mark any corresponding buffers in the buffer cache as stale to prevent log writeback.

  5. 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:

  1. EFIs logged on behalf of space that is no longer occupied

  2. 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 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 length index. Call it again for the free space by block number 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 series.

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 live rebuilds of rmap data, 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 case study.

The proposed patchset is the AG btree repair 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.

See fs/xfs/scrub/agheader_repair.c for more details.

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 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 live quotacheck.

The proposed patchset is the quota repair series.

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 stabilized 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, though not all of the pieces are necessary:

  • The final freeze state is set one higher than SB_FREEZE_COMPLETE to prevent other threads from thawing the filesystem, or other scrub threads from initiating another fscounters freeze.

  • 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.

Historical Sidebar:

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 proposed patchset is the summary counter cleanup 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?

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)" 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", 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:

    1. Move the examination cursor to the point of the inumber keyspace that corresponds to the start of the next AG.

    2. 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.

    3. Unlock the AGI and return to step 1 if there are unexamined AGs in the filesystem.

    4. 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:

    1. Move the examination cursor ahead to the next inode marked as allocated by the inode btree.

    2. 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:

    1. Lock the inode to prevent updates during the scan.

    2. Scan the inode.

    3. While still holding the inode lock, adjust the visited inode cursor (xchk_iscan_mark_visited) to point to this inode.

    4. Unlock and release the inode.

  1. Call xchk_iscan_teardown 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 series. The first user of the new functionality is the online quotacheck 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 an existing 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 and dir iget usage.

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 in numerical order of the addresses of their struct inode object 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 numerical order of the addresses of their struct address_space objects. 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.

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 series.

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 static keys.

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.

Live Updates During a Scan

The code paths of the online fsck scanning code and the hooked 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

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_teardown to finish the scan

This functionality is also a part of the inode scanner series.

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:

    1. The dquot is locked.

    2. A quota reservation is added to the dquot's resource usage. The reservation is recorded in the transaction.

    3. 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:

    1. The dquot is locked again.

    2. Quota usage changes are logged and unused reservation is given back to the dquot.

    3. 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:

    1. Grab and lock the inode.

    2. 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.

    3. Unlock and release the inode.

  3. For each dquot in the system:

    1. Grab and lock the dquot.

    2. 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 series.

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 in-memory array 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 coordinated inode scanner, live update hooks, and an in-memory rmap btree 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:

    1. Create a btree cursor for the in-memory btree.

    2. Use the rmap code to add the record to the in-memory btree.

    3. Use the special commit function 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:

    1. Create a btree cursor for the in-memory btree.

    2. Replay the operation into the in-memory btree.

    3. Use the special commit function 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 reap after rmap btree repair.

  12. Free the xfbtree now that it not needed.

The proposed patchset is the rmap repair 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.

Because file forks can consume as much space as the entire filesystem, repairs cannot be staged in memory, even when a paging scheme is available. Therefore, online repair of file-based metadata createas a temporary file in the XFS filesystem, writes a new structure at the correct offsets into the temporary file, and atomically swaps 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 metadata extents with a temporary file requires the owner field of the block headers to match the file being repaired and not the temporary file. The directory, extended attribute, and symbolic link functions were all modified to allow callers to specify owner numbers explicitly.

There is a downside to the reaping process -- 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.

Historical Sidebar:

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.

This lead to the introduction of temporary file staging.

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 inode locking 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 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 repair temporary files 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:

  1. 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.

  2. 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.

  3. 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.

  4. 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.

  5. 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 new XFS_SB_FEAT_INCOMPAT_LOG_ATOMIC_SWAP log-incompatible feature flag in the superblock protects these new log item records from being replayed on old kernels.

The proposed patchset is the atomic extent swap series.

Sidebar: 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:

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 process the exchange. This is encapsulated in xrep_tempswap_contents for scrub operations. 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,

    1. 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".

    2. Create a deferred block mapping update to unmap map1 from file 1.

    3. Create a deferred block mapping update to unmap map2 from file 2.

    4. Create a deferred block mapping update to map map1 into file 2.

    5. Create a deferred block mapping update to map map2 into file 1.

    6. Log the block, quota, and extent count updates for both files.

    7. Extend the ondisk size of either file if necessary.

    8. Log an extent swap done log item for the extent swap intent log item that was read at the start of step 3.

    9. 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.

    10. 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.

    11. Log a new extent swap intent log item reflecting the advanced state of the work item.

    12. 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.

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 so that the second file will be ready to go as soon as the ILOCK is dropped.

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 builds every block in the new data structure with the owner field of the file being repaired.

After a successful swap operation, the repair operation must reap the old fork blocks by processing each fork mapping through the standard file extent 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.

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:

    1. Compute the position in the summary file that contains a counter that represents this free extent.

    2. Read the counter from the xfile.

    3. 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 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,

    1. 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 series.

Fixing Directories

Fixing directories is difficult with currently available filesystem features, since directory entries are not redundant. 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 file link count fsck code takes care of fixing link counts and moving orphans to the /lost+found directory.

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,

    1. Walk the directory data block to find candidate entries. When an entry is found:

      1. Check the name for problems, and ignore the name if there are.

      2. 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.

Future Work Question: Should repair revalidate the dentry cache when rebuilding a directory?

Answer: Yes, it should.

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 the problem case.

Unfortunately, the current dentry cache design doesn't provide a means to walk every child dentry of a specific directory, which makes this a hard problem. There is no known solution.

The proposed patchset is the directory repair series.

Parent Pointers

A parent pointer is a piece of file metadata that enables a user to locate the file's parent directory without having to traverse the directory tree from the root. Without them, reconstruction of directory trees is hindered in much the same way that the historic lack of reverse space mapping information once hindered reconstruction of filesystem space metadata. The parent pointer feature, however, makes total directory reconstruction possible.

XFS parent pointers include the dirent name and location of the entry within the parent directory. In other words, child files use extended attributes to store pointers to parents in the form (parent_inum, parent_gen, dirent_pos) (dirent_name). 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. Likewise, each parent pointer can be checked by ensuring that the target of each parent pointer is a directory and that it contains a dirent matching the parent pointer. Both online and offline repair can use this strategy.

Note: The ondisk format of parent pointers is not yet finalized.

Historical Sidebar:

Directory parent pointers were first proposed as an XFS feature more than a decade ago by SGI. Each link from a parent directory to a child file is mirrored with an extended attribute in the child that could be used to identify the parent directory. Unfortunately, this early implementation had major shortcomings and was never merged into Linux XFS:

  1. The XFS codebase of the late 2000s did not have the infrastructure to enforce strong referential integrity in the directory tree. It did not guarantee that a change in a forward link would always be followed up with the corresponding change to the reverse links.

  2. Referential integrity was not integrated into offline repair. Checking and repairs were performed on mounted filesystems without taking any kernel or inode locks to coordinate access. It is not clear how this actually worked properly.

  3. The extended attribute did not record the name of the directory entry in the parent, so the SGI 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.

The original parent pointer design was too unstable for something like a file system repair to depend on. Allison Henderson, Chandan Babu, and Catherine Hoang are working on a second implementation that solves all shortcomings of the first. During 2022, Allison introduced log intent items to track physical manipulations of the extended attribute structures. This solves the referential integrity problem by making it possible to commit a dirent update and a parent pointer update in the same transaction. Chandan increased the maximum extent counts of both data and attribute forks, thereby ensuring that the extended attribute structure can grow to handle the maximum hardlink count of any file.

Case Study: Repairing Directories with Parent Pointers

Directory rebuilding uses a coordinated inode scan and a directory entry live update hook as follows:

  1. Set up a temporary directory for generating the new directory structure, an xfblob for storing entry names, and an xfarray for stashing directory updates.

  2. Set up an inode scanner and hook into the directory entry code to receive updates on directory operations.

  3. For each parent pointer found in each file scanned, decide if the parent pointer references the directory of interest. If so:

    1. Stash an addname entry for this dirent in the xfarray for later.

    2. When finished scanning that file, flush the stashed updates to the temporary directory.

  4. For each live directory update received via the hook, decide if the child has already been scanned. If so:

    1. Stash an addname or removename entry for this dirent update in the xfarray for later. We cannot write directly to the temporary directory because hook functions are not allowed to modify filesystem metadata. Instead, we stash updates in the xfarray and rely on the scanner thread to apply the stashed updates to the temporary directory.

  5. When the scan is complete, atomically swap the contents of the temporary directory and the directory being repaired. The temporary directory now contains the damaged directory structure.

  6. Reap the temporary directory.

  7. Update the dirent position field of parent pointers as necessary. This may require the queuing of a substantial number of xattr log intent items.

The proposed patchset is the parent pointers directory repair series.

Unresolved Question: How will repair ensure that the dirent_pos fields match in the reconstructed directory?

Answer: There are a few ways to solve this problem:

  1. 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.

  2. We could 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.

  3. Same as above, but remove the old parent pointer entry and add a new one atomically.

  4. Change the ondisk xattr format to (parent_inum, name) (parent_gen), which would provide the attr name uniqueness that we require, without forcing repair code to update the dirent position. Unfortunately, this requires changes to the xattr code to support attr names as long as 263 bytes.

  5. Change the ondisk xattr format to (parent_inum, hash(name)) (name, parent_gen). If the hash is sufficiently resistant to collisions (e.g. sha256) then this should provide the attr name uniqueness that we require. Names shorter than 247 bytes could be stored directly.

Discussion is ongoing under the parent pointers patch deluge.

Case Study: Repairing Parent Pointers

Online reconstruction of a file's parent pointer information works similarly to directory reconstruction:

  1. Set up a temporary file for generating a new extended attribute structure, an xfblob<xfblob> for storing parent pointer names, and an xfarray for stashing parent pointer updates.

  2. Set up an inode scanner and hook into the directory entry code to receive updates on directory operations.

  3. For each directory entry found in each directory scanned, decide if the dirent references the file of interest. If so:

    1. Stash an addpptr entry for this parent pointer in the xfblob and xfarray for later.

    2. When finished scanning the directory, flush the stashed updates to the temporary directory.

  4. For each live directory update received via the hook, decide if the parent has already been scanned. If so:

    1. Stash an addpptr or removepptr entry for this dirent update in the xfarray for later. We cannot write parent pointers directly to the temporary file because hook functions are not allowed to modify filesystem metadata. Instead, we stash updates in the xfarray and rely on the scanner thread to apply the stashed parent pointer updates to the temporary file.

  5. Copy all non-parent pointer extended attributes to the temporary file.

  6. When the scan is complete, atomically swap the attribute fork of the temporary file and the file being repaired. The temporary file now contains the damaged extended attribute structure.

  7. Reap the temporary file.

The proposed patchset is the parent pointers repair series.

Digression: Offline Checking of Parent Pointers

Examining parent pointers in offline repair works differently because corrupt files are erased long before directory tree connectivity checks are performed. Parent pointer checks are therefore a second pass to be added to the existing connectivity checks:

  1. After the set of surviving files has been established (i.e. phase 6), walk the surviving directories of each AG in the filesystem. This is already performed as part of the connectivity checks.

  2. For each directory entry found, record the name in an xfblob, and store (child_ag_inum, parent_inum, parent_gen, dirent_pos) tuples in a per-AG in-memory slab.

  3. For each AG in the filesystem,

    1. Sort the per-AG tuples in order of child_ag_inum, parent_inum, and dirent_pos.

    2. For each inode in the AG,

      1. Scan the inode for parent pointers. Record the names in a per-file xfblob, and store (parent_inum, parent_gen, dirent_pos) tuples in a per-file slab.

      2. Sort the per-file tuples in order of parent_inum, and dirent_pos.

      3. Position one slab cursor at the start of the inode's records in the per-AG tuple slab. This should be trivial since the per-AG tuples are in child inumber order.

      4. Position a second slab cursor at the start of the per-file tuple slab.

      5. Iterate the two cursors in lockstep, comparing the parent_ino and dirent_pos fields of the records under each cursor.

        1. Tuples in the per-AG list but not the per-file list are missing and need to be written to the inode.

        2. Tuples in the per-file list but not the per-AG list are dangling and need to be removed from the inode.

        3. For tuples in both lists, update the parent_gen and name components of the parent pointer if necessary.

  4. Move on to examining link counts, as we do today.

The proposed patchset is the offline parent pointers repair series.

Rebuilding directories from parent pointers in offline repair is very challenging because it currently uses a single-pass scan of the filesystem during phase 3 to decide which files are corrupt enough to be zapped. This scan would have to be converted into a multi-pass scan:

  1. The first pass of the scan zaps corrupt inodes, forks, and attributes much as it does now. Corrupt directories are noted but not zapped.

  2. The next pass records parent pointers pointing to the directories noted as being corrupt in the first pass. This second pass may have to happen after the phase 4 scan for duplicate blocks, if phase 4 is also capable of zapping directories.

  3. The third pass resets corrupt directories to an empty shortform directory. Free space metadata has not been ensured yet, so repair cannot yet use the directory building code in libxfs.

  4. At the start of phase 6, space metadata have been rebuilt. Use the parent pointer information recorded during step 2 to reconstruct the dirents and add them to the now-empty directories.

This code has not yet been constructed.

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.

Without parent pointers, 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 such a file has a positive link count, the file is an orphan.

With parent pointers, directories can be rebuilt by scanning parent pointers and parent pointers can be rebuilt by scanning directories. This should reduce the incidence of files ending up in /lost+found.

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 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.

Checking Metadata

Recall the phases of fsck work 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:

  1. Filesystem summary counts depend on consistency within the inode indices, the allocation group space btrees, and the realtime volume space information.

  2. 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.

  3. The naming hierarchy depends on consistency within the directory and extended attribute structures. This includes file link counts.

  4. Directories, extended attributes, and file data depend on consistency within the file forks that map directory and extended attribute data to physical storage media.

  5. 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.

  6. Inode records depends on consistency within the inode metadata indices.

  7. Realtime space metadata depend on the inode records and data forks of the realtime metadata inodes.

  8. 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.

  9. 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 scans inodes in parallel. For each inode, groups (f), (e), and (d) are checked, in that order.

  • 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 and the inode scan rebalance series.

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 data dependencies 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.

    1. For each repair item queued by phase 2,

      1. Ask the kernel to repair everything listed in the repair item for a given filesystem object.

      2. Make a note if the kernel made any progress in reducing the number of repairs needed for this object.

      3. 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.

    2. If any repairs were made, jump back to 1a to retry all the phase 2 items.

    3. For each repair item queued by phase 3,

      1. Ask the kernel to repair everything listed in the repair item for a given filesystem object.

      2. Make a note if the kernel made any progress in reducing the number of repairs needed for this object.

      3. 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.

    4. 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, refactoring of the repair data dependency and object tracking, and the 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.

  • Null bytes are not allowed in userspace-visible extended attributes.

  • Null bytes are not allowed in 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 document. When xfs_scrub detects UTF-8 encoding in use on a system, it uses the Unicode normalization form NFD in conjunction with the confusable name detection component of libicu 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, parent pointers can be used to construct file paths from inode numbers for user-friendly reporting.

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.


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.

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.

Vectorized Scrub

As it turns out, the refactoring 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 and userspace 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 atomic extent swap 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.

Future Work Question: How might the filesystem move inode chunks?

Answer: To move inode chunks, Dave Chinner constructed a prototype program 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.

Future Work Question: Can static keys be used to minimize the cost of supporting revoke() on XFS files?

Answer: Yes. Until the first revocation, the bailout code need not be in the call path at all.

The relevant patchsets are the kernel freespace defrag and userspace freespace defrag 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!