linux/fs/xfs/Makefile

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# SPDX-License-Identifier: GPL-2.0
#
# Copyright (c) 2000-2005 Silicon Graphics, Inc.
# All Rights Reserved.
#
ccflags-y += -I $(srctree)/$(src) # needed for trace events
ccflags-y += -I $(srctree)/$(src)/libxfs
obj-$(CONFIG_XFS_FS) += xfs.o
# this one should be compiled first, as the tracing macros can easily blow up
xfs-y += xfs_trace.o
# build the libxfs code first
xfs-y += $(addprefix libxfs/, \
xfs_ag.o \
xfs_alloc.o \
xfs_alloc_btree.o \
xfs_attr.o \
xfs_attr_leaf.o \
xfs_attr_remote.o \
xfs_bit.o \
xfs_bmap.o \
xfs_bmap_btree.o \
xfs_btree.o \
xfs_btree_staging.o \
xfs_da_btree.o \
xfs_defer.o \
xfs_dir2.o \
xfs_dir2_block.o \
xfs_dir2_data.o \
xfs_dir2_leaf.o \
xfs_dir2_node.o \
xfs_dir2_sf.o \
xfs_dquot_buf.o \
xfs_exchmaps.o \
xfs_ialloc.o \
xfs_ialloc_btree.o \
xfs: use a b+tree for the in-core extent list Replace the current linear list and the indirection array for the in-core extent list with a b+tree to avoid the need for larger memory allocations for the indirection array when lots of extents are present. The current extent list implementations leads to heavy pressure on the memory allocator when modifying files with a high extent count, and can lead to high latencies because of that. The replacement is a b+tree with a few quirks. The leaf nodes directly store the extent record in two u64 values. The encoding is a little bit different from the existing in-core extent records so that the start offset and length which are required for lookups can be retreived with simple mask operations. The inner nodes store a 64-bit key containing the start offset in the first half of the node, and the pointers to the next lower level in the second half. In either case we walk the node from the beginninig to the end and do a linear search, as that is more efficient for the low number of cache lines touched during a search (2 for the inner nodes, 4 for the leaf nodes) than a binary search. We store termination markers (zero length for the leaf nodes, an otherwise impossible high bit for the inner nodes) to terminate the key list / records instead of storing a count to use the available cache lines as efficiently as possible. One quirk of the algorithm is that while we normally split a node half and half like usual btree implementations we just spill over entries added at the very end of the list to a new node on its own. This means we get a 100% fill grade for the common cases of bulk insertion when reading an inode into memory, and when only sequentially appending to a file. The downside is a slightly higher chance of splits on the first random insertions. Both insert and removal manually recurse into the lower levels, but the bulk deletion of the whole tree is still implemented as a recursive function call, although one limited by the overall depth and with very little stack usage in every iteration. For the first few extents we dynamically grow the list from a single extent to the next powers of two until we have a first full leaf block and that building the actual tree. The code started out based on the generic lib/btree.c code from Joern Engel based on earlier work from Peter Zijlstra, but has since been rewritten beyond recognition. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Darrick J. Wong <darrick.wong@oracle.com> Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com>
2017-11-03 17:34:46 +00:00
xfs_iext_tree.o \
xfs_inode_fork.o \
xfs_inode_buf.o \
xfs_log_rlimit.o \
xfs: set up per-AG free space reservations One unfortunate quirk of the reference count and reverse mapping btrees -- they can expand in size when blocks are written to *other* allocation groups if, say, one large extent becomes a lot of tiny extents. Since we don't want to start throwing errors in the middle of CoWing, we need to reserve some blocks to handle future expansion. The transaction block reservation counters aren't sufficient here because we have to have a reserve of blocks in every AG, not just somewhere in the filesystem. Therefore, create two per-AG block reservation pools. One feeds the AGFL so that rmapbt expansion always succeeds, and the other feeds all other metadata so that refcountbt expansion never fails. Use the count of how many reserved blocks we need to have on hand to create a virtual reservation in the AG. Through selective clamping of the maximum length of allocation requests and of the length of the longest free extent, we can make it look like there's less free space in the AG unless the reservation owner is asking for blocks. In other words, play some accounting tricks in-core to make sure that we always have blocks available. On the plus side, there's nothing to clean up if we crash, which is contrast to the strategy that the rough draft used (actually removing extents from the freespace btrees). Signed-off-by: Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Dave Chinner <david@fromorbit.com>
2016-09-19 00:30:52 +00:00
xfs_ag_resv.o \
xfs_rmap.o \
xfs_rmap_btree.o \
xfs_refcount.o \
xfs_refcount_btree.o \
xfs_sb.o \
xfs_symlink_remote.o \
xfs_trans_inode.o \
xfs_trans_resv.o \
xfs_types.o \
)
# xfs_rtbitmap is shared with libxfs
xfs-$(CONFIG_XFS_RT) += $(addprefix libxfs/, \
xfs_rtbitmap.o \
)
# highlevel code
xfs-y += xfs_aops.o \
xfs_attr_inactive.o \
xfs_attr_list.o \
xfs_bmap_util.o \
xfs_bio_io.o \
xfs_buf.o \
xfs: test dir/attr hash when loading module Back in the 6.2-rc1 days, Eric Whitney reported a fstests regression in ext4 against generic/454. The cause of this test failure was the unfortunate combination of setting an xattr name containing UTF8 encoded emoji, an xattr hash function that accepted a char pointer with no explicit signedness, signed type extension of those chars to an int, and the 6.2 build tools maintainers deciding to mandate -funsigned-char across the board. As a result, the ondisk extended attribute structure written out by 6.1 and 6.2 were not the same. This discrepancy, in fact, had been noticeable if a filesystem with such an xattr were moved between any two architectures that don't employ the same signedness of a raw "char" declaration. The only reason anyone noticed is that x86 gcc defaults to signed, and no such -funsigned-char update was made to e2fsprogs, so e2fsck immediately started reporting data corruption. After a day and a half of discussing how to handle this use case (xattrs with bit 7 set anywhere in the name) without breaking existing users, Linus merged his own patch and didn't tell the maintainer. None of the ext4 developers realized this until AUTOSEL announced that the commit had been backported to stable. In the end, this problem could have been detected much earlier if there had been any useful tests of hash function(s) in use inside ext4 to make sure that they always produce the same outputs given the same inputs. The XFS dirent/xattr name hash takes a uint8_t*, so I don't think it's vulnerable to this problem. However, let's avoid all this drama by adding our own self test to check that the da hash produces the same outputs for a static pile of inputs on various platforms. This enables us to fix any breakage that may result in a controlled fashion. The buffer and test data are identical to the patches submitted to xfsprogs. Link: https://lore.kernel.org/linux-ext4/Y8bpkm3jA3bDm3eL@debian-BULLSEYE-live-builder-AMD64/ Link: https://lore.kernel.org/linux-xfs/ZBUKCRR7xvIqPrpX@destitution/T/#md38272cc684e2c0d61494435ccbb91f022e8dee4 Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Dave Chinner <dchinner@redhat.com>
2023-03-16 16:31:20 +00:00
xfs_dahash_test.o \
xfs_dir2_readdir.o \
xfs_discard.o \
xfs_error.o \
xfs_exchrange.o \
xfs_export.o \
xfs_extent_busy.o \
xfs_file.o \
xfs_filestream.o \
xfs_fsmap.o \
xfs_fsops.o \
xfs_globals.o \
xfs_health.o \
xfs_icache.o \
xfs_ioctl.o \
xfs_iomap.o \
xfs_iops.o \
xfs_inode.o \
xfs_itable.o \
xfs_iwalk.o \
xfs_message.o \
xfs_mount.o \
xfs_mru_cache.o \
xfs_pwork.o \
xfs_reflink.o \
xfs_stats.o \
xfs_super.o \
xfs_symlink.o \
xfs_sysfs.o \
xfs_trans.o \
xfs_xattr.o
# low-level transaction/log code
xfs-y += xfs_log.o \
xfs_log_cil.o \
xfs_bmap_item.o \
xfs_buf_item.o \
xfs_buf_item_recover.o \
xfs_dquot_item_recover.o \
xfs_exchmaps_item.o \
xfs_extfree_item.o \
xfs: Set up infrastructure for log attribute replay Currently attributes are modified directly across one or more transactions. But they are not logged or replayed in the event of an error. The goal of log attr replay is to enable logging and replaying of attribute operations using the existing delayed operations infrastructure. This will later enable the attributes to become part of larger multi part operations that also must first be recorded to the log. This is mostly of interest in the scheme of parent pointers which would need to maintain an attribute containing parent inode information any time an inode is moved, created, or removed. Parent pointers would then be of interest to any feature that would need to quickly derive an inode path from the mount point. Online scrub, nfs lookups and fs grow or shrink operations are all features that could take advantage of this. This patch adds two new log item types for setting or removing attributes as deferred operations. The xfs_attri_log_item will log an intent to set or remove an attribute. The corresponding xfs_attrd_log_item holds a reference to the xfs_attri_log_item and is freed once the transaction is done. Both log items use a generic xfs_attr_log_format structure that contains the attribute name, value, flags, inode, and an op_flag that indicates if the operations is a set or remove. [dchinner: added extra little bits needed for intent whiteouts] Signed-off-by: Allison Henderson <allison.henderson@oracle.com> Reviewed-by: Chandan Babu R <chandanrlinux@gmail.com> Reviewed-by: Darrick J. Wong <djwong@kernel.org> Signed-off-by: Dave Chinner <david@fromorbit.com>
2022-05-04 02:41:02 +00:00
xfs_attr_item.o \
xfs_icreate_item.o \
xfs_inode_item.o \
xfs_inode_item_recover.o \
xfs_iunlink_item.o \
xfs_refcount_item.o \
xfs_rmap_item.o \
xfs_log_recover.o \
xfs_trans_ail.o \
xfs_trans_buf.o
# optional features
xfs-$(CONFIG_XFS_QUOTA) += xfs_dquot.o \
xfs_dquot_item.o \
xfs_trans_dquot.o \
xfs_qm_syscalls.o \
xfs_qm_bhv.o \
xfs_qm.o \
xfs_quotaops.o
# xfs_rtbitmap is shared with libxfs
xfs-$(CONFIG_XFS_RT) += xfs_rtalloc.o
xfs-$(CONFIG_XFS_POSIX_ACL) += xfs_acl.o
xfs-$(CONFIG_SYSCTL) += xfs_sysctl.o
xfs-$(CONFIG_COMPAT) += xfs_ioctl32.o
xfs-$(CONFIG_EXPORTFS_BLOCK_OPS) += xfs_pnfs.o
# notify failure
ifeq ($(CONFIG_MEMORY_FAILURE),y)
xfs-$(CONFIG_FS_DAX) += xfs_notify_failure.o
endif
xfs: allow queued AG intents to drain before scrubbing When a writer thread executes a chain of log intent items, the AG header buffer locks will cycle during a transaction roll to get from one intent item to the next in a chain. Although scrub takes all AG header buffer locks, this isn't sufficient to guard against scrub checking an AG while that writer thread is in the middle of finishing a chain because there's no higher level locking primitive guarding allocation groups. When there's a collision, cross-referencing between data structures (e.g. rmapbt and refcountbt) yields false corruption events; if repair is running, this results in incorrect repairs, which is catastrophic. Fix this by adding to the perag structure the count of active intents and make scrub wait until it has both AG header buffer locks and the intent counter reaches zero. One quirk of the drain code is that deferred bmap updates also bump and drop the intent counter. A fundamental decision made during the design phase of the reverse mapping feature is that updates to the rmapbt records are always made by the same code that updates the primary metadata. In other words, callers of bmapi functions expect that the bmapi functions will queue deferred rmap updates. Some parts of the reflink code queue deferred refcount (CUI) and bmap (BUI) updates in the same head transaction, but the deferred work manager completely finishes the CUI before the BUI work is started. As a result, the CUI drops the intent count long before the deferred rmap (RUI) update even has a chance to bump the intent count. The only way to keep the intent count elevated between the CUI and RUI is for the BUI to bump the counter until the RUI has been created. A second quirk of the intent drain code is that deferred work items must increment the intent counter as soon as the work item is added to the transaction. When a BUI completes and queues an RUI, the RUI must increment the counter before the BUI decrements it. The only way to accomplish this is to require that the counter be bumped as soon as the deferred work item is created in memory. In the next patches we'll improve on this facility, but this patch provides the basic functionality. Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Dave Chinner <dchinner@redhat.com>
2023-04-12 01:59:58 +00:00
xfs-$(CONFIG_XFS_DRAIN_INTENTS) += xfs_drain.o
xfs: allow scrub to hook metadata updates in other writers Certain types of filesystem metadata can only be checked by scanning every file in the entire filesystem. Specific examples of this include quota counts, file link counts, and reverse mappings of file extents. Directory and parent pointer reconstruction may also fall into this category. File scanning is much trickier than scanning AG metadata because we have to take inode locks in the same order as the rest of [VX]FS, we can't be holding buffer locks when we do that, and scanning the whole filesystem takes time. Earlier versions of the online repair patchset relied heavily on fsfreeze as a means to quiesce the filesystem so that we could take locks in the proper order without worrying about concurrent updates from other writers. Reviewers of those patches opined that freezing the entire fs to check and repair something was not sufficiently better than unmounting to run fsck offline. I don't agree with that 100%, but the message was clear: find a way to repair things that minimizes the quiet period where nobody can write to the filesystem. Generally, building btree indexes online can be split into two phases: a collection phase where we compute the records that will be put into the new btree; and a construction phase, where we construct the physical btree blocks and persist them. While it's simple to hold resource locks for the entirety of the two phases to ensure that the new index is consistent with the rest of the system, we don't need to hold resource locks during the collection phase if we have a means to receive live updates of other work going on elsewhere in the system. The goal of this patch, then, is to enable online fsck to learn about metadata updates going on in other threads while it constructs a shadow copy of the metadata records to verify or correct the real metadata. To minimize the overhead when online fsck isn't running, we use srcu notifiers because they prioritize fast access to the notifier call chain (particularly when the chain is empty) at a cost to configuring notifiers. Online fsck should be relatively infrequent, so this is acceptable. The intended usage model is fairly simple. Code that modifies a metadata structure of interest should declare a xfs_hook_chain structure in some well defined place, and call xfs_hook_call whenever an update happens. Online fsck code should define a struct notifier_block and use xfs_hook_add to attach the block to the chain, along with a function to be called. This function should synchronize with the fsck scanner to update whatever in-memory data the scanner is collecting. When finished, xfs_hook_del removes the notifier from the list and waits for them all to complete. Originally, I selected srcu notifiers over blocking notifiers to implement live hooks because they seemed to have fewer impacts to scalability. The per-call cost of srcu_notifier_call_chain is higher (19ns) than blocking_notifier_ (4ns) in the single threaded case, but blocking notifiers use an rwsem to stabilize the list. Cacheline bouncing for that rwsem is costly to runtime code when there are a lot of CPUs running regular filesystem operations. If there are no hooks installed, this is a total waste of CPU time. Therefore, I stuck with srcu notifiers, despite trading off single threaded performance for multithreaded performance. I also wasn't thrilled with the very high teardown time for srcu notifiers, since the caller has to wait for the next rcu grace period. This can take a long time if there are a lot of CPUs. Then I discovered the jump label implementation of static keys. Jump labels use kernel code patching to replace a branch with a nop sled when the key is disabled. IOWs, they can eliminate the overhead of _call_chain when there are no hooks enabled. This makes blocking notifiers competitive again -- scrub runs faster because teardown of the chain is a lot cheaper, and runtime code only pays the rwsem locking overhead when scrub is actually running. With jump labels enabled, calls to empty notifier chains are elided from the call sites when there are no hooks registered, which means that the overhead is 0.36ns when fsck is not running. This is perfect for most of the architectures that XFS is expected to run on (e.g. x86, powerpc, arm64, s390x, riscv). For architectures that don't support jump labels (e.g. m68k) the runtime overhead of checking the static key is an atomic counter read. This isn't great, but it's still cheaper than taking a shared rwsem. Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Christoph Hellwig <hch@lst.de>
2024-02-22 20:30:45 +00:00
xfs-$(CONFIG_XFS_LIVE_HOOKS) += xfs_hooks.o
xfs-$(CONFIG_XFS_MEMORY_BUFS) += xfs_buf_mem.o
xfs-$(CONFIG_XFS_BTREE_IN_MEM) += libxfs/xfs_btree_mem.o
xfs: allow queued AG intents to drain before scrubbing When a writer thread executes a chain of log intent items, the AG header buffer locks will cycle during a transaction roll to get from one intent item to the next in a chain. Although scrub takes all AG header buffer locks, this isn't sufficient to guard against scrub checking an AG while that writer thread is in the middle of finishing a chain because there's no higher level locking primitive guarding allocation groups. When there's a collision, cross-referencing between data structures (e.g. rmapbt and refcountbt) yields false corruption events; if repair is running, this results in incorrect repairs, which is catastrophic. Fix this by adding to the perag structure the count of active intents and make scrub wait until it has both AG header buffer locks and the intent counter reaches zero. One quirk of the drain code is that deferred bmap updates also bump and drop the intent counter. A fundamental decision made during the design phase of the reverse mapping feature is that updates to the rmapbt records are always made by the same code that updates the primary metadata. In other words, callers of bmapi functions expect that the bmapi functions will queue deferred rmap updates. Some parts of the reflink code queue deferred refcount (CUI) and bmap (BUI) updates in the same head transaction, but the deferred work manager completely finishes the CUI before the BUI work is started. As a result, the CUI drops the intent count long before the deferred rmap (RUI) update even has a chance to bump the intent count. The only way to keep the intent count elevated between the CUI and RUI is for the BUI to bump the counter until the RUI has been created. A second quirk of the intent drain code is that deferred work items must increment the intent counter as soon as the work item is added to the transaction. When a BUI completes and queues an RUI, the RUI must increment the counter before the BUI decrements it. The only way to accomplish this is to require that the counter be bumped as soon as the deferred work item is created in memory. In the next patches we'll improve on this facility, but this patch provides the basic functionality. Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Dave Chinner <dchinner@redhat.com>
2023-04-12 01:59:58 +00:00
# online scrub/repair
ifeq ($(CONFIG_XFS_ONLINE_SCRUB),y)
# Tracepoints like to blow up, so build that before everything else
xfs-y += $(addprefix scrub/, \
trace.o \
agb_bitmap.o \
agheader.o \
alloc.o \
attr.o \
bitmap.o \
bmap.o \
btree.o \
common.o \
dabtree.o \
dir.o \
fscounters.o \
health.o \
ialloc.o \
inode.o \
2024-02-22 20:30:45 +00:00
iscan.o \
listxattr.o \
nlinks.o \
parent.o \
readdir.o \
refcount.o \
rmap.o \
scrub.o \
symlink.o \
xfarray.o \
xfile.o \
)
xfs-$(CONFIG_XFS_ONLINE_SCRUB_STATS) += scrub/stats.o
xfs-$(CONFIG_XFS_RT) += $(addprefix scrub/, \
rtbitmap.o \
rtsummary.o \
)
xfs-$(CONFIG_XFS_QUOTA) += $(addprefix scrub/, \
dqiterate.o \
quota.o \
quotacheck.o \
)
# online repair
ifeq ($(CONFIG_XFS_ONLINE_REPAIR),y)
xfs-y += $(addprefix scrub/, \
agheader_repair.o \
alloc_repair.o \
attr_repair.o \
bmap_repair.o \
cow_repair.o \
dir_repair.o \
findparent.o \
fscounters_repair.o \
ialloc_repair.o \
inode_repair.o \
xfs: implement block reservation accounting for btrees we're staging Create a new xrep_newbt structure to encapsulate a fake root for creating a staged btree cursor as well as to track all the blocks that we need to reserve in order to build that btree. As for the particular choice of lowspace thresholds and btree block slack factors -- at this point one could say that the thresholds in online repair come from bulkload_estimate_ag_slack in xfs_repair[1]. But that's not the entire story, since the offline btree rebuilding code in xfs_repair was merged as a retroport of the online btree code in this patchset! Before xfs_btree_staging.[ch] came along, xfs_repair determined the slack factor (aka the number of slots to leave unfilled in each new btree block) via open-coded logic in repair/phase5.c[2]. At that point the slack factors were arbitrary quantities per btree. The rmapbt automatically left 10 slots free; everything else left zero. That had a noticeable effect on performance straight after mounting because adding records to /any/ btree would result in splits. A few years ago when this patch was first written, Dave and I decided that repair should generate btree blocks that were 75% full unless space was tight, in which case it should try to fill the blocks to nearly full. We defined tight as ~10% free to avoid repair failures but settled on 3/32 (~9%) to avoid div64. IOWs, we mostly pulled the thresholds out of thin air. We've been QAing with those geometry numbers ever since. ;) Link: https://git.kernel.org/pub/scm/fs/xfs/xfsprogs-dev.git/tree/repair/bulkload.c?h=v6.5.0#n114 Link: https://git.kernel.org/pub/scm/fs/xfs/xfsprogs-dev.git/tree/repair/phase5.c?h=v4.19.0#n1349 Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Dave Chinner <dchinner@redhat.com>
2023-12-07 02:40:59 +00:00
newbt.o \
nlinks_repair.o \
orphanage.o \
parent_repair.o \
xfs: define an in-memory btree for storing refcount bag info during repairs Create a new in-memory btree type so that we can store refcount bag info in a much more memory-efficient and performant format. Recall that the refcount recordset regenerator computes the new recordset from browsing the rmap records. Let's say that the rmap records are: {agbno: 10, length: 40, ...} {agbno: 11, length: 3, ...} {agbno: 12, length: 20, ...} {agbno: 15, length: 1, ...} It is convenient to have a data structure that could quickly tell us the refcount for an arbitrary agbno without wasting memory. An array or a list could do that pretty easily. List suck because of the pointer overhead. xfarrays are a lot more compact, but we want to minimize sparse holes in the xfarray to constrain memory usage. Maintaining any kind of record order isn't needed for correctness, so I created the "rcbag", which is shorthand for an unordered list of (excerpted) reverse mappings. So we add the first rmap to the rcbag, and it looks like: 0: {agbno: 10, length: 40} The refcount for agbno 10 is 1. Then we move on to block 11, so we add the second rmap: 0: {agbno: 10, length: 40} 1: {agbno: 11, length: 3} The refcount for agbno 11 is 2. We move on to block 12, so we add the third: 0: {agbno: 10, length: 40} 1: {agbno: 11, length: 3} 2: {agbno: 12, length: 20} The refcount for agbno 12 and 13 is 3. We move on to block 14, and remove the second rmap: 0: {agbno: 10, length: 40} 1: NULL 2: {agbno: 12, length: 20} The refcount for agbno 14 is 2. We move on to block 15, and add the last rmap. But we don't care where it is and we don't want to expand the array so we put it in slot 1: 0: {agbno: 10, length: 40} 1: {agbno: 15, length: 1} 2: {agbno: 12, length: 20} The refcount for block 15 is 3. Notice how order doesn't matter in this list? That's why repair uses an unordered list, or "bag". The data structure is not a set because it does not guarantee uniqueness. That said, adding and removing specific items is now an O(n) operation because we have no idea where that item might be in the list. Overall, the runtime is O(n^2) which is bad. I realized that I could easily refactor the btree code and reimplement the refcount bag with an xfbtree. Adding and removing is now O(log2 n), so the runtime is at least O(n log2 n), which is much faster. In the end, the rcbag becomes a sorted list, but that's merely a detail of the implementation. The repair code doesn't care. (Note: That horrible xfs_db bmap_inflate command can be used to exercise this sort of rcbag insanity by cranking up refcounts quickly.) Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Christoph Hellwig <hch@lst.de>
2024-02-22 20:43:40 +00:00
rcbag_btree.o \
rcbag.o \
reap.o \
refcount_repair.o \
repair.o \
rmap_repair.o \
symlink_repair.o \
tempfile.o \
xfblob.o \
)
xfs-$(CONFIG_XFS_RT) += $(addprefix scrub/, \
rtbitmap_repair.o \
rtsummary_repair.o \
)
xfs-$(CONFIG_XFS_QUOTA) += $(addprefix scrub/, \
quota_repair.o \
quotacheck_repair.o \
)
endif
endif