| /linux-6.12.1/drivers/md/bcache/ |
| D | journal.h | 9 * never spans two buckets. This means (not implemented yet) we can resize the
15 * We also keep some things in the journal header that are logically part of the
20 * rewritten when we want to move/wear level the main journal.
22 * Currently, we don't journal BTREE_REPLACE operations - this will hopefully be
25 * moving gc we work around it by flushing the btree to disk before updating the
35 * We track this by maintaining a refcount for every open journal entry, in a
38 * zero, we pop it off - thus, the size of the fifo tells us the number of open
41 * We take a refcount on a journal entry when we add some keys to a journal
42 * entry that we're going to insert (held by struct btree_op), and then when we
43 * insert those keys into the btree the btree write we're setting up takes a
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| D | bset.h | 17 * We use two different functions for validating bkeys, bch_ptr_invalid and
27 * them on disk, just unnecessary work - so we filter them out when resorting
30 * We can't filter out stale keys when we're resorting, because garbage
32 * unless we're rewriting the btree node those stale keys still exist on disk.
34 * We also implement functions here for removing some number of sectors from the
44 * There could be many of them on disk, but we never allow there to be more than
45 * 4 in memory - we lazily resort as needed.
47 * We implement code here for creating and maintaining auxiliary search trees
48 * (described below) for searching an individial bset, and on top of that we
62 * Since keys are variable length, we can't use a binary search on a bset - we
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| /linux-6.12.1/fs/btrfs/ |
| D | space-info.c | 27 * 1) space_info. This is the ultimate arbiter of how much space we can use.
30 * reservations we care about total_bytes - SUM(space_info->bytes_) when
35 * metadata reservation we have. You can see the comment in the block_rsv
39 * 3) btrfs_calc*_size. These are the worst case calculations we used based
40 * on the number of items we will want to modify. We have one for changing
41 * items, and one for inserting new items. Generally we use these helpers to
47 * We call into either btrfs_reserve_data_bytes() or
48 * btrfs_reserve_metadata_bytes(), depending on which we're looking for, with
49 * num_bytes we want to reserve.
66 * Assume we are unable to simply make the reservation because we do not have
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| D | delalloc-space.c | 23 * We call into btrfs_reserve_data_bytes() for the user request bytes that
24 * they wish to write. We make this reservation and add it to
25 * space_info->bytes_may_use. We set EXTENT_DELALLOC on the inode io_tree
27 * make a real allocation if we are pre-allocating or doing O_DIRECT.
30 * At writepages()/prealloc/O_DIRECT time we will call into
31 * btrfs_reserve_extent() for some part or all of this range of bytes. We
35 * may allocate a smaller on disk extent than we previously reserved.
46 * This is the simplest case, we haven't completed our operation and we know
47 * how much we reserved, we can simply call
60 * We keep track of two things on a per inode bases
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| D | direct-io.c | 62 * We're concerned with the entire range that we're going to be in lock_extent_direct()
63 * doing DIO to, so we need to make sure there's no ordered in lock_extent_direct()
70 * We need to make sure there are no buffered pages in this in lock_extent_direct()
71 * range either, we could have raced between the invalidate in in lock_extent_direct()
90 * If we are doing a DIO read and the ordered extent we in lock_extent_direct()
91 * found is for a buffered write, we can not wait for it in lock_extent_direct()
92 * to complete and retry, because if we do so we can in lock_extent_direct()
99 * range and this range started (we unlock the ranges in lock_extent_direct()
112 * We could trigger writeback for this range (and wait in lock_extent_direct()
118 * ordered dio extent we created before but did not have in lock_extent_direct()
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| D | fiemap.c | 30 * - Cache the next entry to be emitted to the fiemap buffer, so that we can
35 * buffer is memory mapped to the fiemap target file, we don't deadlock
36 * during btrfs_page_mkwrite(). This is because during fiemap we are locking
40 * if the fiemap buffer is memory mapped to the file we are running fiemap
53 * the next file extent item we must search for in the inode's subvolume
59 * This matches struct fiemap_extent_info::fi_mapped_extents, we use it
61 * fiemap_fill_next_extent() because we buffer ready fiemap entries at
62 * the @entries array, and we want to stop as soon as we hit the max
86 * Ignore 1 (reached max entries) because we keep track of that in flush_fiemap_cache()
102 * And only when we fails to merge, cached one will be submitted as
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| /linux-6.12.1/fs/xfs/ |
| D | xfs_log_cil.c | 23 * recover, so we don't allow failure here. Also, we allocate in a context that
24 * we don't want to be issuing transactions from, so we need to tell the
27 * We don't reserve any space for the ticket - we are going to steal whatever
28 * space we require from transactions as they commit. To ensure we reserve all
29 * the space required, we need to set the current reservation of the ticket to
30 * zero so that we know to steal the initial transaction overhead from the
42 * set the current reservation to zero so we know to steal the basic in xlog_cil_ticket_alloc()
62 * We can't rely on just the log item being in the CIL, we have to check
80 * current sequence, we're in a new checkpoint. in xlog_item_in_current_chkpt()
140 * We're in the middle of switching cil contexts. Reset the in xlog_cil_push_pcp_aggregate()
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| D | xfs_log_priv.h | 74 * By covering, we mean changing the h_tail_lsn in the last on-disk
83 * might include space beyond the EOF. So if we just push the EOF a
91 * system is idle. We need two dummy transaction because the h_tail_lsn
103 * we are done covering previous transactions.
104 * NEED -- logging has occurred and we need a dummy transaction
106 * DONE -- we were in the NEED state and have committed a dummy
108 * NEED2 -- we detected that a dummy transaction has gone to the
110 * DONE2 -- we committed a dummy transaction when in the NEED2 state.
112 * There are two places where we switch states:
114 * 1.) In xfs_sync, when we detect an idle log and are in NEED or NEED2.
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| D | xfs_log.c | 77 * We need to make sure the buffer pointer returned is naturally aligned for the
78 * biggest basic data type we put into it. We have already accounted for this
81 * However, this padding does not get written into the log, and hence we have to
86 * We also add space for the xlog_op_header that describes this region in the
87 * log. This prepends the data region we return to the caller to copy their data
89 * is not 8 byte aligned, we have to be careful to ensure that we align the
90 * start of the buffer such that the region we return to the call is 8 byte
171 * we have overrun available reservation space, return 0. The memory barrier
289 * path. Hence any lock will be globally hot if we take it unconditionally on
292 * As tickets are only ever moved on and off head->waiters under head->lock, we
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| D | xfs_log_recover.c | 78 * Pass log block 0 since we don't have an addr yet, buffer will be in xlog_alloc_buffer()
88 * We do log I/O in units of log sectors (a power-of-2 multiple of the in xlog_alloc_buffer()
89 * basic block size), so we round up the requested size to accommodate in xlog_alloc_buffer()
97 * blocks (sector size 1). But otherwise we extend the buffer by one in xlog_alloc_buffer()
249 * h_fs_uuid is null, we assume this log was last mounted in xlog_header_check_mount()
328 * range of basic blocks we'll be examining. If that fails, in xlog_find_verify_cycle()
329 * try a smaller size. We need to be able to read at least in xlog_find_verify_cycle()
330 * a log sector, or we're out of luck. in xlog_find_verify_cycle()
385 * a good log record. Therefore, we subtract one to get the block number
387 * of blocks we would have read on a previous read. This happens when the
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| /linux-6.12.1/net/ipv4/ |
| D | tcp_vegas.c | 15 * o We do not change the loss detection or recovery mechanisms of
19 * only every-other RTT during slow start, we increase during
22 * we use the rate at which ACKs come back as the "actual"
24 * o To speed convergence to the right rate, we set the cwnd
25 * to achieve the right ("actual") rate when we exit slow start.
26 * o To filter out the noise caused by delayed ACKs, we use the
55 /* There are several situations when we must "re-start" Vegas:
60 * o when we send a packet and there is no outstanding
63 * In these circumstances we cannot do a Vegas calculation at the
64 * end of the first RTT, because any calculation we do is using
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| /linux-6.12.1/arch/powerpc/mm/nohash/ |
| D | tlb_low_64e.S | 91 /* We need _PAGE_PRESENT and _PAGE_ACCESSED set */
93 /* We do the user/kernel test for the PID here along with the RW test
95 /* We pre-test some combination of permissions to avoid double
98 * We move the ESR:ST bit into the position of _PAGE_BAP_SW in the PTE
103 * writeable, we will take a new fault later, but that should be
106 * We also move ESR_ST in _PAGE_DIRTY position
109 * MAS1 is preset for all we need except for TID that needs to
137 * We are entered with:
176 /* Now we build the MAS:
219 /* We need to check if it was an instruction miss */
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| /linux-6.12.1/fs/xfs/scrub/ |
| D | agb_bitmap.c | 19 * We know that the btree query_all function starts at the left edge and walks
20 * towards the right edge of the tree. Therefore, we know that we can walk up
22 * to the first record/key in that block, we haven't seen this block before;
23 * and therefore we need to remember that we saw this block in the btree.
32 * the first btree record, we'll observe that bc_levels[0].ptr == 1, so we
33 * record that we saw block 1. Then we observe that bc_levels[1].ptr == 1, so
34 * we record block 4. The list is [1, 4].
36 * For the second btree record, we see that bc_levels[0].ptr == 2, so we exit
39 * For the 101st btree record, we've moved onto leaf block 2. Now
40 * bc_levels[0].ptr == 1 again, so we record that we saw block 2. We see that
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| D | alloc_repair.c | 48 * AG. Therefore, we can recreate the free extent records in an AG by looking
60 * walking the rmapbt records, we create a second bitmap @not_allocbt_blocks to
72 * The OWN_AG bitmap itself isn't needed after this point, so what we really do
83 * written to the new btree indices. We reconstruct both bnobt and cntbt at
84 * the same time since we've already done all the work.
86 * We use the prefix 'xrep_abt' here because we regenerate both free space
118 * Next block we anticipate seeing in the rmap records. If the next
119 * rmap record is greater than next_agbno, we have found unused space.
126 /* Longest free extent we found in the AG. */
138 * Make sure the busy extent list is clear because we can't put extents in xrep_setup_ag_allocbt()
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| D | fscounters.c | 31 * The basics of filesystem summary counter checking are that we iterate the
34 * Then we compare what we computed against the in-core counters.
37 * While we /could/ freeze the filesystem and scramble around the AGs counting
38 * the free blocks, in practice we prefer not do that for a scan because
39 * freezing is costly. To get around this, we added a per-cpu counter of the
40 * delalloc reservations so that we can rotor around the AGs relatively
41 * quickly, and we allow the counts to be slightly off because we're not taking
42 * any locks while we do this.
44 * So the first thing we do is warm up the buffer cache in the setup routine by
47 * structures as quickly as it can. We snapshot the percpu counters before and
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| /linux-6.12.1/Documentation/filesystems/xfs/ |
| D | xfs-delayed-logging-design.rst | 15 We begin with an overview of transactions in XFS, followed by describing how
16 transaction reservations are structured and accounted, and then move into how we
18 reservations bounds. At this point we need to explain how relogging works. With
113 individual modification is atomic, the chain is *not atomic*. If we crash half
140 complete, we can explicitly tag a transaction as synchronous. This will trigger
145 throughput to the IO latency limitations of the underlying storage. Instead, we
161 available to write the modification into the journal before we start making
164 log in the worst case. This means that if we are modifying a btree in the
165 transaction, we have to reserve enough space to record a full leaf-to-root split
166 of the btree. As such, the reservations are quite complex because we have to
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| /linux-6.12.1/fs/bcachefs/ |
| D | bset.h | 21 * We use two different functions for validating bkeys, bkey_invalid and
27 * them on disk, just unnecessary work - so we filter them out when resorting
30 * We can't filter out stale keys when we're resorting, because garbage
32 * unless we're rewriting the btree node those stale keys still exist on disk.
34 * We also implement functions here for removing some number of sectors from the
44 * There could be many of them on disk, but we never allow there to be more than
45 * 4 in memory - we lazily resort as needed.
47 * We implement code here for creating and maintaining auxiliary search trees
48 * (described below) for searching an individial bset, and on top of that we
62 * Since keys are variable length, we can't use a binary search on a bset - we
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| /linux-6.12.1/drivers/misc/vmw_vmci/ |
| D | vmci_route.c | 33 * which comes from the VMX, so we know it is coming from a in vmci_route()
36 * To avoid inconsistencies, test these once. We will test in vmci_route()
37 * them again when we do the actual send to ensure that we do in vmci_route()
49 * If this message already came from a guest then we in vmci_route()
57 * We must be acting as a guest in order to send to in vmci_route()
63 /* And we cannot send if the source is the host context. */ in vmci_route()
71 * then they probably mean ANY, in which case we in vmci_route()
87 * If it is not from a guest but we are acting as a in vmci_route()
88 * guest, then we need to send it down to the host. in vmci_route()
89 * Note that if we are also acting as a host then this in vmci_route()
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| /linux-6.12.1/drivers/usb/dwc2/ |
| D | hcd_queue.c | 32 /* If we get a NAK, wait this long before retrying */
121 * @num_bits: The number of bits we need per period we want to reserve
123 * @interval: How often we need to be scheduled for the reservation this
127 * the interval or we return failure right away.
128 * @only_one_period: Normally we'll allow picking a start anywhere within the
129 * first interval, since we can still make all repetition
131 * here then we'll return failure if we can't fit within
134 * The idea here is that we want to schedule time for repeating events that all
139 * To keep things "simple", we'll represent our schedule with a bitmap that
141 * but does mean that we need to handle things specially (and non-ideally) if
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| /linux-6.12.1/Documentation/filesystems/ |
| D | directory-locking.rst | 10 When taking the i_rwsem on multiple non-directory objects, we
11 always acquire the locks in order by increasing address. We'll call
22 * lock the directory we are accessing (shared)
26 * lock the directory we are accessing (exclusive)
73 in its own right; it may happen as part of lookup. We speak of the
74 operations on directory trees, but we obviously do not have the full
75 picture of those - especially for network filesystems. What we have
77 Trees grow as we do operations; memory pressure prunes them. Normally
78 that's not a problem, but there is a nasty twist - what should we do
83 possibility that directory we see in one place gets moved by the server
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| D | idmappings.rst | 23 on, we will always prefix ids with ``u`` or ``k`` to make it clear whether
24 we're talking about an id in the upper or lower idmapset.
42 that make it easier to understand how we can translate between idmappings. For
43 example, we know that the inverse idmapping is an order isomorphism as well::
49 Given that we are dealing with order isomorphisms plus the fact that we're
50 dealing with subsets we can embed idmappings into each other, i.e. we can
51 sensibly translate between different idmappings. For example, assume we've been
61 Because we're dealing with order isomorphic subsets it is meaningful to ask
64 mapping ``k11000`` up to ``u1000``. Afterwards, we can map ``u1000`` down using
69 If we were given the same task for the following three idmappings::
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| /linux-6.12.1/drivers/gpu/drm/i915/ |
| D | i915_request.c | 72 * We could extend the life of a context to beyond that of all in i915_fence_get_timeline_name()
74 * or we just give them a false name. Since in i915_fence_get_timeline_name()
129 * freed when the slab cache itself is freed, and so we would get in i915_fence_release()
138 * We do not hold a reference to the engine here and so have to be in i915_fence_release()
139 * very careful in what rq->engine we poke. The virtual engine is in i915_fence_release()
140 * referenced via the rq->context and we released that ref during in i915_fence_release()
141 * i915_request_retire(), ergo we must not dereference a virtual in i915_fence_release()
142 * engine here. Not that we would want to, as the only consumer of in i915_fence_release()
147 * we know that it will have been processed by the HW and will in i915_fence_release()
153 * power-of-two we assume that rq->engine may still be a virtual in i915_fence_release()
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| /linux-6.12.1/drivers/gpu/drm/i915/gt/ |
| D | intel_execlists_submission.c | 24 * shouldn't we just need a set of those per engine command streamer? This is
35 * Regarding the creation of contexts, we have:
43 * like before) we need:
50 * more complex, because we don't know at creation time which engine is going
51 * to use them. To handle this, we have implemented a deferred creation of LR
55 * gets populated for a given engine once we receive an execbuffer. If later
56 * on we receive another execbuffer ioctl for the same context but a different
57 * engine, we allocate/populate a new ringbuffer and context backing object and
61 * only allowed with the render ring, we can allocate & populate them right
96 * we use a NULL second context) or the first two requests have unique IDs.
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| /linux-6.12.1/kernel/irq/ |
| D | spurious.c | 26 * We wait here for a poller to finish.
28 * If the poll runs on this CPU, then we yell loudly and return
32 * We wait until the poller is done and then recheck disabled and
33 * action (about to be disabled). Only if it's still active, we return
86 * All handlers must agree on IRQF_SHARED, so we test just the in try_one_irq()
209 * We need to take desc->lock here. note_interrupt() is called in __report_bad_irq()
210 * w/o desc->lock held, but IRQ_PROGRESS set. We might race in __report_bad_irq()
244 /* We didn't actually handle the IRQ - see if it was misrouted? */ in try_misrouted_irq()
249 * But for 'irqfixup == 2' we also do it for handled interrupts if in try_misrouted_irq()
260 * Since we don't get the descriptor lock, "action" can in try_misrouted_irq()
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| /linux-6.12.1/arch/openrisc/mm/ |
| D | fault.c | 59 * We fault-in kernel-space virtual memory on-demand. The in do_page_fault()
62 * NOTE! We MUST NOT take any locks for this case. We may in do_page_fault()
68 * mappings we don't have to walk all processes pgdirs and in do_page_fault()
69 * add the high mappings all at once. Instead we do it as they in do_page_fault()
82 /* If exceptions were enabled, we can reenable them here */ in do_page_fault()
100 * If we're in an interrupt or have no user in do_page_fault()
101 * context, we must not take the fault.. in do_page_fault()
125 * we get page-aligned addresses so we can only check in do_page_fault()
126 * if we're within a page from usp, but that might be in do_page_fault()
137 * Ok, we have a good vm_area for this memory access, so in do_page_fault()
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