A "zombie file" is a term I just made up to describe what happens when
you remove a file that is currently open.
To match POSIX, the opened file handle should still be available for
reading/writing, even though the file doesn't really exist in the
filesystem anymore.
We don't have inodes, which makes this a bit more complicated, but this
is where scratch files are handy again. By creating a scratch file when
we remove an opened file, we preserve the mid slot for the file's
sprout/shrub. We also mark the opened file as desync, so the existing
orphan reclaimation circuitry kicks in when the last file handle is
closed.
Really the only difference between zombie files and desync files is what
happens when you call lfsr_file_sync:
- Desynced lfsr_file_sync => Become synced, broadcast file state.
- Zombied lfsr_file_sync => Return ENOENT, you can't sync a zombie.
This _is_ a bit different from POSIX, where sync on a removed file
returns 0. I considered returning 0 in this case, but with all the extra
behavior around sync/desync state, I figured returning ENOENT was
clearer at indicating to the user sync is no longer possible.
Worst case, ENOENT is not returned from sync for any other reason, so
users can always treat ENOENT and 0 as the same in higher layers. The
zombie file is already desynced, so close will never error.
---
Implementation wise, zombies get a bit crazy.
Fortunately they add little extra code, but they make up for it by
adding extra subtlety. Zombie files introduce a ton of corner cases, now
even directories can have zombied shrubs.
This means more tests.
- Seemingly unrelated operations need to be able to remove scratch files
(mkdir, rename, etc).
- UNCREAT state needs to be broadcasted in seemingly unrelated
operations (mkdir, rename, etc).
- Zombied files need to be copied over during seemingly unrelated rename
operations.
- And I'm sure more corner cases I'm already forgetting.
One interesting tweak that simplifies things that's worth mentioning is
the change to the implicitly file mid updates on rm in lfsr_mdir_commit.
For non-reg files, an rm attr causes lfsr_mdir_commit to increment the
mid to the next mid in the mtree. This is the correct behavior for dirs,
traversals, etc.
Previously, reg files were a special case that marks the mid as -1. But
by changing this to also increment the mid, as well as set the zombie
flag, upper layers can broadcast zombie changes by simply creating a new
file and then deleting the old file in the same commit.
This seems to Just Work^TM, and avoids needing to do additional state
broadcasting in upper layers, which gets tricky since we may not know
exactly what the new mid is post-mdir-commit.
Downside: The order matters, we need to create the new file first. This
violates the normal delete-then-insert order we use elsewhere to avoid
overflow issues. This isn't that bad here, since we increment by at
most 1. But it is something to be wary of...
Still, this is much better than any other option I can think of right
now.
---
Uh, ignore the test_fscratch_rename* tests for now. I somehow forgot
file renaming was not yet implemented...
"Scratch files" are a new file type added to solve the zero-sized
file problem. Though they have a few other uses that may be quite
valuable.
The "zero-sized file problem" is a common surprise for users, where what
seems like a simple file create+write operation:
lfs_file_open(&lfs, &file, "hi",
LFS_O_WRONLY | LFS_O_CREAT | LFS_O_EXCL);
lfs_file_write(&lfs, &file, "hello!", strlen("hello!"));
lfs_file_close(&lfs, &file);
Can end up create a zero-sized file under powerloss, breaking user
assumptions and their code.
The tricky thing is that this is actually correct behavior as defined by
POSIX. `open` with O_CREAT creats a file entry immediately, which is
initially zero-sized. And the fact that power can be lost between `open`
and `close` isn't really avoidable.
But this is a common enough footgun that it's probably worth deviating
from POSIX here.
But how to avoid zero-sized files exactly? First thought: Delay the file
creation until sync/close, tracking uncreated files in-device until
then. This solves the problem and avoids any intermediary state if we
lose power, but came with a number of headaches:
1. Since we delay file creation, we don't immediately write the filename
to disk on open. This implies we need to keep the filename allocated
in RAM until the first sync/close call.
The requirement to keep the filename allocated for new files until
first sync/close could be added to open, and with the option to call
sync immediately to save the filename (and accept the risk of
zero-sized files), I don't think it would be _that_ bad of an API.
But it would still be pretty bad. Extra bad because 1. there's no
way to warn on misuse at compile-time, 2. use-after-free bugs have a
tendency to go unnoticed annoyingly often, 3. it's a regression from
the previous API, and 4. who the heck reads the more-or-less same
`open` documentation for every filesystem they adopt.
2. Without an allocated mid, tracking files internally gets a lot
harder. The best option I could think of was to keep the opened-file
linked-list sorted by mid + (in-device) file name.
This did not feel like a great solutiona and was going to add more
code cost.
3. Handling mdir splits containing uncreated files adds another
headache. Complicated lfsr_mdir_estimate further as it needs to
decide in which mdir the uncreated files will end up, and potentially
split on a filename that isn't even created yet.
4. Since the number of uncreated files can be potentially unbounded, you
can't prevent an mdir from filling up with only uncreated files. On
disk this ends up looking like an "empty" mdir, which need specially
handling in littlefs to reclaim after powerloss.
Support for empty mdirs -- the orphaned mdir scan -- was already
added earlier. We already scan each mdir to build gstate, so it
doesn't really add much cost.
Notice that last bullet point? We already scan each mdir during mount.
Why not, instead of scanning for orphaned mdirs, scan for orphaned
files?
So this leads to the idea of "scratch files". Instead of actually
delaying file creation, fake it. Create a scratch file during open, and
on the first sync/close, convert it to a regular file. If we lose power,
scan for scratch files during mount, and remove them on first write.
Some tradeoffs:
1. The orphan scan for scratch files is a bit more expensive than for
mdirs on storage with large block sizes. We need to look at each file
entry vs just each mdir, which pushed the runtime up to O(BlogB) vs
O(B).
Though if you also consider large mtrees, the worst case is still
O(nlogn).
2. Creating intermediate scratch files adds another commit to file
creation.
This is probably not a big issue for flash, but may be more of a
concern on devices with large prog sizes.
3. Scratch files complicate unrelated mkdir/rename/etc code a bit, since
we need to consider what happens when the dest is a scratch file.
But the end result is simple. And simple is good. Both for
implementation headaches, and code size. Even if the on-disk state is
conceptually more complicated.
You may have noticed these scratch files are basically isomorphic to
just setting an "uncreated" flag on the file, and that's true. There may
have been a simpler route to end up with the design, but hey, as long as
it works.
As a plus, scratch files present a solution for a couple other things:
1. Removing an open file can become a scratch file until closed.
2. Scratch files can be used as temporary files. Open a file with
O_DESYNC and never call sync and you have yourself a temporary file.
Maybe in the future we should add O_TMPFILE to avoid the need for
unique filenames, but that is low priority.
This should, in theory, be a transparent change for users
(https://xkcd.com/1172).
The motivation for this change:
1. Basically everyone uses O_RDONLY=0, O_WRONLY=1, O_RDWR=2, so
deviating from this ad-hoc standard risks surprising POSIX-familiar
users, though may confused POSIX-unfamiliar users.
But for the latter, we really shouldn't allow them to fall into the
trap that O_RDONLY | O_WRONLY == O_RDWR, because this will not work
on basically any other POSIX-like system.
2. I realized one benefit of the POSIX encoding is that it reserves the
value 3. Maybe this could be useful in the future?
Being able to create a file that neither readable nor writable isn't
all that useful...
Also, if you really think about the literal meaning of O_RDONLY |
O_WRONLY, these are negations. So O_RDONLY | O_WRONLY means you can only
write and only read? That sounds like an oxymoron.
Of course no one should be relying on these exact values, but these are
embedded systems! Someone somewhere is going to hack something together
that expects these to be their historically expected value. And we
shouldn't make things any harder for them unless there's a good reason.
While the multi per-type linked-lists were cool and could save RAM in
some structs (at the cost of RAM in the lfs_t struct), this is simpler,
and simpler is good.
The motivation to revert:
1. I noticed most file types have some sort of flags: files,
traversals (future), (not dirs but maybe in the future). These flags
can be merged with the type field to give us typed mdirs at almost
no RAM cost.
2. Using a single linked-list makes it cheaper to add more file types,
which may be useful for managing bookmarks (differently) and scratch
files.
This comes at a runtime cost, since all scans look at all opened
structs, but we really, _really_ don't care about a constant non-IO
runtime cost.
There are code benefits, since we don't need nested iterators to access
all opened mdirs, but also some code cost when we want to filter by
type. As expected stack took a small hit. Humorously, the struct savings
in lfs_t perfectly canceled out the struct hit to lfsr_dir_t:
code stack structs
before: 32992 2968 1080
after: 33004 (+0.0%) 2976 (+0.3%) 1080 (+0.0%)
This readds lfsr_ftree_t, however this time its not involved in the file
staging, has no operations of its own, and really just acts as a
namespace for the file's bnull/bsprout/bptr/bshrub/btree struct.
I think this is a good way to organize things.
Code impact is also minimal:
code stack
before: 32874 2952
after: 32984 (+0.3%) 2968 (+0.5%)
One less struct to worry about, and less code/stack pressure from
passing around multiple pointers.
There were some naming collisions:
- lfsr_ftree_size -> lfsr_file_bsize
- lfsr_ftree_read -> lfsr_file_read_
- lfsr_ftree_flush -> lfsr_file_flush_
I'm not sure this should be the final result. There are definitely some
rough spots, the hacky "pseudo-file" in lfsr_traversal_t for example.
Having a name specific to file btrees was also useful for naming/
documenting things...
But the code savings are hard to shake a stick at:
code stack
before: 33260 3024
after: 32874 (-1.2%) 2952 (-2.4%)
Direct block pointers are turning out to be a bit of an awkward file
representation for littlefs. Thanks to shrubs, direct block pointers
really don't offer that much in terms of disk savings.
Direct bptrs save ~40 B:
direct bptr: 1 attr + 1 bptr
40 B + 24 B = 64 B
indirect bshrub: 2 attr + 1 trunk + 1 bptr
2*40 B + 10 B + 24 B = 114 B
δ = +40 B (+78.1%)
Which is nice, but not really significant on disk. Their original
motivation was to avoid the cost of a btree root node for one block
files. But this can now be avoided with bshrubs, which also generalizes
to other few-block files.
I can see the argument for carving out a special case for entirely
inlined files. +~40B may be a significant cost there. But I'm just not
seeing the value for bptrs.
But direct bptrs exist as a natural extension of littlefs's design.
Files can have:
1. nothing, null data,
2. a data entry (bptr/bsprout)
3. a bshrub/btree of data entries (bptr/bsprout)
Prohibiting direct bptrs, would be a bit strange, and a future version
of littlefs may find direct bptrs useful. Say, for example, a version
that doesn't support bshrubs, suddenly bptrs become more valuable.
So this is a compromise:
1. Support reading of bptrs, this is not that much extra work on top of
supporting bsprouts. Though we do need to be aware of them in the
block allocator.
2. Convert bptrs to bshrubs/btrees on first write.
3. Ignore any extra bptr metadata, becksums, cids, etc. These add an
additional attr which complicates things.
Downside: We may lose out on potential erased-state when writing to
files created on a different device that uses bptrs. Upside: Simpler
code and a bit of code savings.
code stack
before: 33260 3024
after: 33136 (-0.4%) 3000 (-0.8%)
Ok, maybe not that much code savings...
This is an attempt to simplify things a bit by moving more logic into
the ftree layer, instead of spreading things around between the
bshrub/bsprout functions.
Now, functionality is organized into high-level ftree operations and
low-level shrub/sprout operations, which only care about the inlined
portion of the shrub/sprout. No more lfsr_bshrub_commit/
lfsr_bshrub_commit__ which were mostly unrelated.
This also adds a lfsr_shrub_t type, which, by taking advantage of the
unused write-related rbyd fields to store the shrub estimate, has the
same size as lfsr_rbyd_t, but can still be casted to an rbyd/btree for
use in readonly rbyd/btree functions.
I considered merging shrub/sprout esimate and shrub/sprout compact into
some sort of ftree_estimate/compact, but it's not obvious what the
benefit would be, so leaving that on the table for now.
---
One nice change is our staging copies are now at the ftree level
(ftree.u and ftree.u_, maybe not the best names, but this is what I've
been using for unions where the name doesn't really matter, god I want
unnamed unions). This simplifies staging, and avoids staging issues
where the underlying type changes.
---
A bit unrelated, but necessary to integrate lfsr_ftree_traverse, a
generalized lfsr_tinfo_t type for all traversal functions was added
(adopted from lfsr_traversal_t really). This is a straightforward tagged
union with relevant traversal types.
The benefit of a generalized tinfo type is better chance we can just
pass the tinfo pointer through multiple layers.
Code changes:
code stack
before: 33368 2984
after: 33260 (-0.3%) 3024 (+1.3%)
Note, I think if we ever add file snapshots for idempotent errors again,
I don't think adding mdir/next back into the ftree is the best way to
structure this.
Instead, adding a separate linked-list for tracking bshrubs would work
without adding redundant mdir copies to the ftree struct.
Fortunately, in our current version, we don't need to track on-stack
ftrees. Actually, we don't make on-stack ftree copies at all...
While convenient, file.size is redundant info. Redundant info always
has the risk of falling out-of-sync, creating difficult to find bugs.
This was made especially apparent with dropping file-level idempotent
errors, which make possible file states quite a bit more complex (we've
given up on fully reverting errors, but we don't want errors to make the
filesystem inconsistent).
Replacing file.size with an inlinable function that derives the file
size removes this risk without too much cost. As a plus, lfsr_file_t is
one word smaller:
code stack lfsr_file_t
before: 33286 2968 112
after: 33278 (-0.0%) 2976 (+0.3%) 108 (-3.6%)
A recent change, motivated by user feedback, was to delay write buffer
flushes as much as possible. Before, littlefs would always flush the
buffer during lfs_file_seek, but now, buffer flushes can be delayed all
the way to lfsr_file_read, or even skipped entirely thanks to bypassing
reads.
This is all fine and dandy, except it's easy to imagine a use case where
a user might really not want a _write_ error to pop out of a _read_
call.
With this new behavior, avoiding this situation is impossible.
So enters a function common to other filesystems: lfsr_file_flush.
However it's value is quite a bit different here. Unlike flush in other
filesystems, this flush does not necessarily make data accessible on
disk. It only writes to the pending file snapshot, which is not
accessible until lfsr_file_sync.
This makes flush a function with a rather narrow scope in littlefs
(pretty much just preventing write errors in read), but since we had
already implemented this function for internal plumbing, it adds _very_
little cost.
I'm more concerned about potential user confusion around sync vs flush.
Curiously, exposing lfsr_file_flush actually _saved_ code size for some
reason. Not sure what would make that happen:
code stack
before: 33544 3072
flush: 33536 (-0.0%) 3072 (+0.0%)
flush+O_FLUSH: 33548 (+0.0%) 3072 (+0.0%)
The motivation for this comes from the observation that many users call
sync on every file write. Much more than I expected. I think one reason
is in embedded systems it's common to just write structs to disk, either
the whole file or to a log.
O_SYNC exists in POSIX/Lunix/etc, so it makes sense to provide in
littlefs. In theory it's just one extra function call, and may even save
in total application cost (though we don't measure this) by reducing the
number of function calls at the application-level.
---
Unfortunately in-practice turned out to be quite a bit different than
in-theory... The main culprit being the improved guarantees around error
atomicity...
The ideal guarantee is that if there is an error during a write, the
entire write operation is reverted. Combining this with O_SYNC means we
need to hold a copy of the origin file state all thwe way through our
sync call. This got a bit messy...
The annoying part isn't even the functionality! Our system of tracking
btree/bshrub snapshots is quite robust! The problems were entirely with:
1. Figuring out how the heck to avoid clobbering the old file buffer
state.
2. Figuring out how the internal APIs should work while passing around a
bunch of staging state.
For 1., fortunately, thanks to bypassing writes, and some careful
pointer manipulation, we can void buffer clobbing. And for 2. just some
internal API work was needed. Internally all syncs end up in
lfsr_ftree_sync, though this feels a bit clumsy since the functionality
is not really ftree related...
Unfortunately, all of this added up to quite a bit more code cost than
I had hoped. In theory, adding some sort of LFS_CERAMIC/LFS_GLASS modes
that relax error atomicity for code size could help with most of this?
But it needs some thought:
code stack
before: 33324 3072
after: 33544 (+0.7%) 3072 (+0.0%)
None of the available options sit well with me.
Worst case writes states after an error:
1. Maintain on-stack snapshots for entire write operation:
on-disk: abcdefghijklmnopqrstuvwxyz
write: JKLMN
error!
on-disk: abcdefghiJKlmnopqrstuvwxyz
2. Maintain on-stack snapshots for lfsr_ftree_carve:
on-disk: abcdefghijklmnopqrstuvwxyz
write: JKLMN
error!
on-disk: abcdefghijklmnopqrstuvwxyz
3. Don't maintain on-stack snapshots, rely on btree/bshrub atomicity:
on-disk: abcdefghijklmnopqrstuvwxyz
write: JKLMN
error!
on-disk: abcdstuvwxyz
Something else to consider, the on-stack snapshots increase pressure on
the available shrub_size, which must include all tracked bshrubs in the
mdir, and currently doesn't deduplicate more than checking for identical
trunks. In effect, shrubs are limited to ~shrub_size/3, which isn't
great...
Since we can't get rid of the extra shrub cost when atomic carve
operations, I'm going to revert this, since we might as well just track
all file operations and provide a fully atomic API... Element of least
surprise and all thath...
But this revert may itself be reverted in the future.
Maybe we should provide some sort of LFS_LESSATOMIC flag to allow opt-in
to non-atomic file writes for code/stack savings?
This is an attempt to reduce the overhead of on-stack snapshots during
writes, by relaxing the gaurantees provided by lfsr_file_write during
errors.
Before, thanks to the on-stack snapshot, file writes could revert to the
previous state if an error occurred. Now, on-stack snapshots are limited
to lfsr_ftree_carve, so only the state change in lfsr_ftree_carve is
reverted.
This should behave relatively predictably, since lfsr_ftree_flush calls
lfsr_ftree_carve in a normal order. If an error occurs, some, none, or
all of the data is actually written. For truncate/fruncate, there is
only one call to carve, so these remain atomic.
It's tempting to want to push this lower. If you push the atomic
operations down to the btree/bshrub level, individual btree/bshrub commits
are already atomic, so tracking on-stack bshrubs could be dropped
completely. But failures at the sub-carve level get weird! Thanks to
block crystallization and the use of order-statistic operations, the
resulting file can be quite unpredictable. For example:
on-disk: abcdefghijklmnopqrstuvwxyz
write: JKLMN
error!
on-disk: abcdstuvwxyz
With atomic carves, worst case is something like this:
on-disk: abcdefghijklmnopqrstuvwxyz
write: JKLMN
error!
on-disk: abcdefghiJKlmnopqrstuvwxyz
Thanks to file-level snapshotting, the original file can still be
recovered (and is on-disk until sync is called), but the
unpredictability is concerning.
Alternatively, if we had btree range removals, lfsr_ftree_carve could be
entirely atomic (and more efficient).
Unfortunately, btree range removals still seem like they would be
difficult to implement, and will probably be out of scope for some time.
The added code cost will also likely outweigh any savings from dropping
on-stack bshrub tracking. Still, this is probably worth looking into in
the future.
Code changes:
code stack
before: 33356 3072
after: 33006 (-1.0%) 2992 (-2.6%)
Desynchronized files are a new concept intended to capture some useful
quirks of the previous multiple-open-file behavior.
This adds:
- LFS_O_DESYNC - Mark a file as desync during open
- lfsr_file_desync - Mark a file as desync whenever
- lfsr_file_sync - Mark a file as NOT desync, and sync the file
Desynced files:
1. Don't recieve updates from writes to other file handles. This makes
desynced files act as a sort of snapshot of the file at the time it
was marked desync.
2. Don't call lfsr_file_sync on close. Unless lfsr_file_sync is
explicitly called, changes to desynced files are not reflected on
disk and not broadcasted to other file handles.
A side-effect of 2., is that this gives you a quick way to abort a file
write. Marking a file as desync and then closing the file will never
error.
Additionally, if an error occurs during a write operation, the file is
implicitly marked as desync. This provides graceful write aborting in
unlikely error cases. This has actually always been a feature in
littlefs, it was just named differently and didn't have an optional
recovery mode.
Since littlefs actually has to do more work to keep files in sync, the
desync feature is quite cheap:
code stack
before: 33324 3072
after: 33360 (+0.1%) 3072 (+0.0%)
Now that bshrubs are limited to files, we might as well lean into it.
Instead of relying on extra traversals through our attr-list, we now
rely on bsprouts/bshrubs always existing in the opened mdir list. This
is a much more robust way to deduplicate bsprout/bshrub compaction
operations. The previous implementation already had a bug since
bsprout/bshrub calculation was forgotten from lfsr_mdir_estimate_.
This is a direct tradeoff of code and RAM. Though note the previous
impl was incomplete and would likely need a bit more code:
code stack
before: 33184 2944
after: 32980 (-0.6%) 3000 (+1.9%)
This is intended to enable lfsr_file_read to use the buffer as well.
This adds LFS_F_UNFLUSHED and internal lfsr_file_flush to manage buffer
flushing. This also results in a nice reorganization of lfsr_file_sync.
Of course, small file caching proves to be a big pain again, with
several subtle corner cases in truncate/fruncate that need to set the
LFS_F_UNFLUSHED flag so we fix small files in lfsr_file_sync.
Our crystallization threshold doesn't really describe the bounds of an
object, and I think it's a bit easier to think of it as a threshold for
block compaction.
Heck I've already been calling this the crystallization threshold all
over the code base.
An important change is this bumps the value by 1 bytes, so
crystal_thresh now describes the smallest size of a block our write
strategy will attempt to write.
Heuristically:
- data >= crystal_thresh => compacted into blocks
- data < crystal_thresh => stored as fragments
Much like the erased-state checksums in our rbyds (ecksums), these
block-level erased-state checksums (becksums) allow us to detect failed
progs to erased parts of a block and are key to achieving efficient
incremental write performance with large blocks and frequent power
cycles/open-close cycles.
These are also key to achieving _reasonable_ write performance for
simple writes (linear, non-overwriting), since littlefs now relies
solely on becksums to efficiently append to blocks.
Though I suppose the previous block staging logic used with the CTZ
skip-list could be brought back to make becksums optional and avoid
btree lookups during simple writes (we do a _lot_ of btree
lookups)... I'll leave this open as a future optimization...
Unlike in-rbyd ecksums, becksums need to be stored out-of-band so our
data blocks only contain raw data. Since they are optional, an
additional tag in the file's btree makes sense.
Becksums are relatively simple, but they bring some challenges:
1. Adding becksums to file btrees is the first case we have for multiple
struct tags per btree id.
This isn't too complicated a problem, but requires some new internal
btree APIs.
Looking forward, which I probably shouldn't be doing this often,
multiple struct tags will also be useful for parity and content ids
as a part of data redundancy and data deduplication, though I think
it's uncontroversial to consider this both heavier-weight features...
2. Becksums only work if unfilled blocks are aligned to the prog_size.
This is the whole point of crystal_size -- to provide temporary
storage for unaligned writes -- but actually aligning the block
during writes turns out to be a bit tricky without a bunch of
unecesssary btree lookups (we already do too many btree lookups!).
The current implementation here discards the pcache to force
alignment, taking advantage of the requirement that
cache_size >= prog_size, but this is corrupting our block checksums.
Code cost:
code stack
before: 31248 2792
after: 32060 (+2.5%) 2864 (+2.5%)
Also lfsr_ftree_flush needs work. I'm usually open to gotos in C when
they improve internal logic, but even for me, the multiple goto jumps
from every left-neighbor lookup into the block writing loop is a bit
much...
Looking forward, bptr checksums provide an easy mechanism to validate
data residing in blocks. This extends the merkle-tree-like nature of the
filesystem all the way down to the data level, and is common in other
COW filesystems.
Two interesting things to note:
1. We don't actually check data-level checksums yet, but we do calculate
data-level checksums unconditionally.
Writing checksums is easy, but validating checksums is a bit more
tricky. This is made a bit harder for littlefs, since we can't hold
an entire block of data in RAM, so we have to choose between separate
bus transactions for checksum + data reads, or extremely expensive
overreads every read.
Note this already exists at the metadata-level, the separate bus
transactions for rbyd fetch + rbyd lookup means we _are_ susceptible
to a very small window where bit errors can get through.
But anyways, writing checksums is easy. And has basically no cost
since we are already processing the data for our write. So we might
as well write the data-level checksums at all times, even if we
aren't validating at the data-level.
2. To make bptr checksums work cheaply we need an additional cksize
field to indicate how much data is checksummed.
This field seems redundant when we already have the bptr's data size,
but if we didn't have this field, we would be forced to recalculate
the checksum every time a block is sliced. This would be
unreasonable.
The immutable cksize field does mean we may be checksumming more data
than we need to when validating, but we should be avoiding small
block slices anyways for storage cost reasons.
This does add some stack cost because our bptr struct is larger now:
code stack
before: 31200 2768
after: 31272 (+0.2%) 2800 (+1.1%)
Also:
- Renamed GSTATE -> GDELTA for gdelta tags. GSTATE tags added as
separate in-device flags. The GSTATE tags were already serving
this dual purpose.
- Renamed BSHRUB* -> SHRUB when the tag is not necessarily operating
on a file bshrub.
- Renamed TRUNK -> BSHRUB
The tag encoding space now has a couple funky holes:
- 0x0005 - Hole for aligning config tags.
I guess this could be used for OCOMPATFLAGS in the future?
- 0x0203 - Hole so that ORPHAN can be a 1-bit difference from REG. This
could be after BOOKMARK, but having a bit to differentiate littlefs
specific file types (BOOKMARK, ORPHAN) from normal file types (REG,
DIR) is nice.
I guess this could be used for SYMLINK if we ever want symlinks in the
future?
- 0x0314-0x0318 - Hole so that the mdir related tags (MROOT, MDIR,
MTREE) are nicely aligned.
This is probably a good place for file-related tags to go in the
future (BECKSUM, CID, COMPR), but we only have two slots, so will
probably run out pretty quickly.
- 0x3028 - Hole so that all btree related tags (BTREE, BRANCH, MTREE)
share a common lower bit-pattern.
I guess this could be used for MSHRUB if we ever want mshrubs in the
future?
lfsr_ftree_t acts as a sort of proto-file type, holding enough
information for file reads/writes if the relevant mdir is known.
This lets low-level file write operations operate on a copy of the
proto-file without needing to copy the relevant mdir, file stuff, etc.
To make this work, lfsr_mdir_commit also needs to stage any bshrubs in
the attr-list, since these may not be in our opened file list, but this
is a good thing to handle implicitly anyways. We should only ever have
one untracked bshrub being operated on (multithreaded support would be a
whole other can of worms).
Unfortunately the extra machinery in lfsr_mdir_commit, and the fact that
passing two pointers around instead of one adds quite a bit of code,
means this comes with a code cost. But the tradeoff for stack cost and
no risk of stack pointers in our opened file list makes this probably
worth it:
code stack
before: 31584 2824
after: 31760 (+0.6%) 2776 (-1.7%)
This ended up being much less of a simplification than I hoped it would.
It's still easier/more efficient to revert to a relocation in most cases
when dropping in an mdir split, and the small gain from simplifying how
drops/commits interact is overshadowed by the code duplication necessary
to separate lfsr_mdir_drop out from lfsr_mdir_commit:
code stack
before: 30952 2528
after: 31280 (+1.1%) 2648 (+4.7%)
Still, this does at least simplify the logical corner cases (we don't
need to abort commits when droppable anymore), and lfsr_mdir_drop is
ultimately necessary for supporting lazy file creation.
Also having a fix-orphans step during mount allows other littlefs
implementations the option to create orphanned mdirs without compat
issues. So this ends up the more flexible approach.
It _might_ be worth having both eager mdir drops and an explicit
lfsr_mdir_drop for lazy file creation in the future, but I doubt this
will end up worth the code duplication...
---
Oh right, I forgot to actually describe this change.
This trades eager mdir drops:
1. Drop mdirs from the mtree immediately as soon as their weight goes
to zero.
For lazy mdir drops:
1. Drop mdirs from the mtree in a second commit.
2. Scan and drop orphaned mdirs on the first write after mount.
This sounds very similar to the previous "deorphan" scan, which risked
an extreme performance cost during mount, but it should be noted this
orphan scan only needs to touch every mdir once. This makes it no worse
than the overhead of actually mounting the filesystem.
We can also keep an eye out for orphaned mdirs when we mount, so no
extra scan is needed unless there was an unlucky powerloss.
Eager mdir dropping sounds simpler, but thanks to deferred commits
introduces some subtle complexity around aborting commits that would
drop an mdir to zero. Remember commits are viewable on-disk as soon as a
commit completes.
In _theory_, lazy mdir drops simplify the logic around committing to
mdirs.
Though the real kicker is that lazy mdir drops are required for lazy file
creation.
The current idea for lazy file creation involves tracking mid-less
opened-but-not-yet-created files. These files can have bshrubs, so they
need space on an mdir somewhere. But they aren't actually created yet,
so they don't have an mid.
This is fine (though it's probably going to be tricky) as long as we
allocate an mid on file sync, but there is always a risk of losing power
with mdirs that contain only RAM-backed files. Fortunately, no-mids
means no orphaned files, but it does mean orphaned mdirs with no synced
contents.
Long story short, lazy mdir drops are currently a necessary evil, and
logical simplification, that unfortunately comes with some cost.
The idea here is to revert moving redund blocks into lfsr_rbyd_t, and
instead just keep a redundant copy of the rbyd blocks in the redund
blocks in lfsr_mdir_t.
Surprisingly, extra overhead in lfsr_mdir_t ended up with worse stack
usage than extra overhead in lfsr_rbyd_t. I guess we end up allocated
more mdirs than rbyds, which makes a bit of sense given how complicated
lfsr_mdir_commit is:
code stack structs
redund union: 30976 2496 1072
redund in rbyd: 30948 (-0.1%) 2528 (+1.3%) 1100 (+2.6%)
redund in mdir: 31000 (+0.1%) 2536 (+1.6%) 1092 (+1.8%)
The mdir option does seem to improve struct overhead, but this hasn't
been a reliable measurement since it doesn't take into account how many
of each struct is allocated.
Given that the mdir option is inferior in both code and stack cost, and
requires more care to keep the rbyd/redund blocks in sync, I think I'm
going to revert this for now but keep the commit in the commit history
since it's an interesting comparison.
This simplifies dependent structs with redundancy, mainly lfsr_mdir_t,
at a significant RAM cost:
code stack structs
before: 30976 2496 1072
after: 30948 (-0.1%) 2528 (+1.3%) 1100 (+2.6%)
Which, to be honest, is not as bad as I thought it would be. Though it
is still pretty bad for no new features.
The motivation for this change:
1. The organization of the previous lfsr_mdir_t struct was a bit hacky
and relied on exact padding so the redund block array and rbyd block
lined up at the right offset.
2. The previous organization prevented theoretical "read-only rbyd
structs" that could omit write-related fields, e.g. eoff and cksum.
This idea is currently unused.
3. The current mdir=level-1, btree/data=level-0 redund design makes this
RAM tradeoff pretty bad, but in theory higher btree redund levels
would need the extra redund blocks in the rbyd struct anyways.
Still, the RAM impact to the current default configuration means this
should probably be reverted...
- Renamed mdir->u.m to mdir->u.mdir.
- Prefer mdir->u.rbyd.* where possible.
- Changed file/dir mdirs to be stored directly, requiring a cast to
lfsr_openedmdir_t to enroll in the opened mdir list.
This is just a useful type to have to make the code a bit more
readable.
This doesn't affect the code that much, except we are making more
on-stack copies of mptrs since the mdir doesn't technically contain
a mutable mptr. Maybe this should change?
code stack
before: 30768 2496
after: 30776 (+0.0%) 2504 (+0.3%)
Unfortunately, waiting to evict shrubs until mdir compaction does not
work because we only have a single pcache. When we evict a bshrub we
need a pcache for writing the new btree root, but if we do this during
mdir compaction, our pcache is already busy handling the mdir
compaction. We can't do a separate pass for bshrub eviction, since this
would require tracking an unbounded number of new btree roots.
In the previous shrub design, we meticulously tracked the compacted
shrub estimate in RAM, determining exactly how the estimate would change
as a part of shrub carve operations.
This worked, but was fragile. It was easy for the shrub estimate to
diverge from the actual value, and required quite a bit of extra code to
maintain. Since the use cases for bshrubs is growing a bit, I didn't
want to return to this design.
So here's a new approach based on emulating btree compacts/splits inside
the shrubs:
1. When a bshrub is fetched, scan the bshrub and calculate a compaction
estimate. Store this.
2. On every commit, find the upper bound of new data being progged, and
keep track of estimate + progged. We can at least get this relatively
easily from commit attr lists. We can't get the amount deleted, which
is the problem.
3. When estimate + progged exceeds shrub_size, scan the bshrub again and
recalculate the estimate.
4. If estimate exceeds the shrub_size/2, evict the bshrub, converting it
into a btree.
As you may note, this is very close to how our btree compacts/splits
work, but emulated. In particular, evictions/splits occur at
(shrub_size/block_size)/2 in order to avoid runaway costs when the
bshrub/btree gets close to full.
Benefits:
- This eviction heuristic is very robust. Calculating the amount progged
from the attr list is relatively cheap and easy, and any divergence
should be fixed when we recalculate the estimate.
- The runtime cost is relatively small, amortized O(log n) which is
the existing runtime to commit to rbyds.
Downsides:
- Just like btree splits, evictions force our bshrub to be ~1/2 full on
average. This combined with the 2x cost for mdir pairs, the 2x cost
for mdirs being ~1/2 full on average, and the need for both a synced
and unsynced copy of file bshrubs brings our file bshrub's overhead up
to ~16x, which is getting quite high...
Anyways, bshrubs now work, and the new file topology is passing testing.
An unfortunate surprise is the jump in stack cost. This seems to come from
moving the lfsr_btree_flush logic into the hot-path that includes bshrub
commit + mdir commit + all the mtree logic. Previously the separate of
btree/shrub commits meant that the more complex block/btree/crystal logic
was on a separate path from the mdir commit logic:
code stack lfsr_file_t
before bshrubs: 31840 2072 120
after bshrubs: 30756 (-3.5%) 2448 (+15.4%) 104 (-15.4%)
I _think_ the reality is not actually as bad as measured, most of these
flush/carve/commit functions calculate some work and then commit it in
seperate steps. In theory GCC's shrinkwrapping optimizations should
limit the stack to only what we need as we finish different
calculations, but our current stack measurement scripts just add
together the whole frames, so any per-call stack optimizations get
missed...
As a part of the general redesign of files, all files, not just small
files, can inline some data directly in the metadata log. Originally,
this was a single piece of inlined data or an inlined tree (shrub) that
effectively acted as an overlay over the block/btree data.
This is now changed so that when we have a block/btree, the root of the
btree is inlined. In effect making a full btree a sort of extended
shrub.
I'm currently calling this a "geoxylic btree", since that seems to be a
somewhat related botanical term. Geoxylic btrees have, at least on
paper, a number of benefits:
- There is a single lookup path instead of two, this simplifies code a
bit and decreases lookup costs.
- One data structure instead of two also means lfsr_file_t requires
less RAM, since all of the on-disk variants can go into one big union.
Though I'm not sure this is very significant vs stack/buffer costs.
- The write path is much simpler and has less duplication (it was
difficult to deduplicate the shrub/btree code because of how the
shrub goes through the mdir).
In this redesign, lfsr_btree_commit_ leaves root attrs uncommitted,
allowing lfsr_bshrub_commit to finish the job via lfsr_mdir_commit.
- We don't need to maintain a shrub estimate, we just lazily evict trees
during mdir compaction. This has a side-effect of allowing shrubs to
temporarily grow larger than shrub_size before eviction.
NOTE THIS (fundamentally?) DOESN'T WORK
- There is no awkwardly high overhead for small btrees. The btree root
for two-block files should be able to comfortably fit in the shrub
portion of the btree, for example.
- It may be possible to also make the mtree geoxylic, which should
reduce storage overhead of small mtrees and make better use of the
mroot.
All of this being said, things aren't working yet. Shrub eviction during
compaction runs into a problem with a single pcache -- how do we write
the new btrees without dropping the compaction pcache? We can't evict
btrees in a separate pass becauce their number is unbounded...
This did not turn out to be useful, mainly because type-agnostic
inlining requires unnecessary encoding/decoding and risks a higher RAM
allocation than is really needed. It's better to just reserve a bit in
the weight field and allow higher-level operations to use
operation-specific unions.
code stack
before: 31580 2072
after: 31160 (-1.3%) 2072 (+0.0%)
This is based on how bench.py/bench_runners have actually been used in
practice. The main changes have been to make the output of bench.py more
readibly consumable by plot.py/plotmpl.py without needing a bunch of
hacky intermediary scripts.
Now instead of a single per-bench BENCH_START/BENCH_STOP, benches can
have multiple named BENCH_START/BENCH_STOP invocations to measure
multiple things in one run:
BENCH_START("fetch", i, STEP);
lfsr_rbyd_fetch(&lfs, &rbyd_, rbyd.block, CFG->block_size) => 0;
BENCH_STOP("fetch");
Benches can also now report explicit results, for non-io measurements:
BENCH_RESULT("usage", i, STEP, rbyd.eoff);
The extra iter/size parameters to BENCH_START/BENCH_RESULT also allow
some extra information to be calculated post-bench. This infomation gets
tagged with an extra bench_agg field to help organize results in
plot.py/plotmpl.py:
- bench_meas=<meas>+amor, bench_agg=raw - amortized results
- bench_meas=<meas>+div, bench_agg=raw - per-byte results
- bench_meas=<meas>+avg, bench_agg=avg - average over BENCH_SEED
- bench_meas=<meas>+min, bench_agg=min - minimum over BENCH_SEED
- bench_meas=<meas>+max, bench_agg=max - maximum over BENCH_SEED
---
Also removed all bench.tomls for now. This may seem counterproductive in
a commit to improve benchmarking, but I'm not sure there's actual value
to keeping bench cases committed in tree.
These were alway quick to fall out of date (at the time of this commit
most of the low-level bench.tomls, rbyd, btree, etc, no longer
compiled), and most benchmarks were one-off collections of scripts/data
with results too large/cumbersome to commit and keep updated in tree.
I think the better way to approach benchmarking is a seperate repo
(multiple repos?) with all related scripts/state/code and results
committed into a hopefully reproducible snapshot. Keeping the
bench.tomls in that repo makes more sense in this model.
There may be some value to having benchmarks in CI in the future, but
for that to make sense they would need to actually fail on performance
regression. How to do that isn't so clear. Anyways we can always address
this in the future rather than now.
This gives the mtree a dedicated type, with direct mptrs (single mdirs)
being stored decoded, instead of encoding into leb128s. This avoids
encoding/decoding in some cases.
This change is currently a net downgrade, but only because we still have
all of the inlined btree code. Eventually this inlined btree code should
be removed:
code stack
before: 31316 2064
after: 31480 (+0.5%) 2072 (+0.4%)
Also tweaked the tests to no longer test dropping the mtree down to
zero size. Thanks to root bookmarks, we never actually do this, and it
simplifies lfsr_mdir_commit to not support this.
- Ripped out outdated file-data representation. We don't need this.
- Changed lfsr_data_add/read/cmp to just assert when data is
concatenated data. Theoretically this is possible to implement, but
it's complicated and we never use it, so all it is is a waste of
code size...
- Added implicitly zero-filled hole representation, though this isn't
adopted in the code yet.
- Added lfsr_data_truncate/fruncate, these are really useful for
shrub/tree carving/coalescing.
---
New lfsr_data_t encoding, sign(size) indicates if the data is
on-disk/in-device, and a mode field indicates how in-device data should
be parsed:
sign(size)=1 => on-disk:
.---+---+---+---. .....
|1| size | ..'' ''..
+---+---+---+---+ : : :
| block ------+->| ..:|
+---+---+---+---+ | |......( )::::::|
| off -------' |:::' : |
'---+---+---+---' :' : :
''.. :.''
'''''
sign(size)=0, mode=0 => in-device buffer:
.---+---+---+---. .---+---+---+---.
|0| size | .>| data... |
+---+---+---+---+ | ' . '
|m=0| | | ' . '
+---+---+---+---+ | ' '
| ptr -------' ' '
'---+---+---+---' '---+---+---+---'
sign(size)=0, mode=1 => hole
.---+---+---+---.
|0| size |
+---+---+---+---+
|m=1| |
+---+ +
| |
'---+---+---+---'
sign(size)=0, mode=2 => inlined
.---+---+---+---.
|0| size |
+---+---+---+---+
|m=2| inlined d |
+---+ +
| ata... |
'---+---+---+---'
sign(size)=0, mode=3 => concatenated datas:
.---+---+---+---. .---+---+---+---.
|0| size | .>| data |
+---+---+---+---+ | + +
|m=3| c | | | | |
+---+---+---+---+ | + +
| ptr -------' | |
'---+---+---+---' +---+---+---+---+
| data |
+ +
| |
+ +
| |
+---+---+---+---+
' . '
' . '
' . '
' '
' '
'---+---+---+---'
---
Code/RAM changes:
code stack
before: 31952 2056
after: 31396 (-1.7%) 2064 (+0.4%)
I think the increased RAM cost is due to lfsr_data_add/truncate/fruncate
passing lfsr_data_t around by value, and GCC not being able to optimize
this very well since it's 3 words. I think most move optimizations stop
after 2-words...
The biggest change here is the breaking up of the FLAGS config into
RFLAGS/WFLAGS/OFLAGS. This is directly inspired by, and honestly not
much more than a renaming, of the compat/ro_compat/incompat flags found
in Linux/Unix/POSIX filesystems.
I think these were first introduced in ext2? But I need to do a bit more
research on that.
RFLAGS/WFLAGS/OFLAGS provide a much more flexible, and extensible,
feature flag mechanism than the previous minor version bumps.
The (re)naming of these flags is intended to make their requirements
more clear. In order to do the relevant operation, you must understand
every flag set in the relevant flag:
- RFLAGS / incompat flags - All flags must be understood to read the
filesystem, if not understood the only possible behavior is to fail.
- WFLAGS / ro-compat flags - All flags must be understood to write to the
filesystem, if not understood the filesystem may be mounted read-only.
- OFLAGS / compat flags - Optional flags, if not understood the relevant
flag must be cleared before the filesystem can be written to, but other
than that these flags can mostly be ignored.
Some hypothetical littlefs examples:
- RFLAGS / incompat flags - Transparent compression
Is this the same as a major disk-version break? Yes kinda? An
implementation that doesn't understand compression can't read the
filesystem.
On the other hand, it's useful to have a filesystem that can read both
compressed and uncompressed variants.
- WFLAGS / ro-compat flags - Closed block-map
The idea behind a closed block-map (currently planned), is that
littlefs maintains in global space a complete mapping of all blocks in
use by the filesystem.
For such a mapping to remain consistent means that if you write to the
filesystem you must understand the closed block-map. Or in other
words, if you don't understand the closed block-map you must not write
to the filesystem.
Reading, on the other hand, can ignore many such write-related
auxiliary features, so the filesystem can still be read from.
- OFLAGS / compat flags - Global checksums
Global checksums (currently planned) are extra checksums attached to
each mdir that when combined self-validate the filesystem.
But if you don't understand global checksums, you can still read and
write the filesystem without them. The only catch is that when you write
to the filesystem, you may end up invalidating the global checksum.
Clearing the global checksum bit in the OFLAGS is a cheap way to
signal that the global checksum is no longer valid, allowing you to
still write to the filesystem without this optional feature.
Other tweaks to note:
- Renamed BLOCKLIMIT/DISKLIMIT -> BLOCKSIZE/BLOCKCOUNT
Note these are still the _actual_ block_size/block_count minus 1. The
subtle difference here was the original reason for the name change,
but after working with it for a bit, I just don't think new, otherwise
unused, names are worth it.
The minus 1 stays, however, since it avoids overflow issues at
extreme boundaries of powers of 2.
- Introduces STAGLIMIT/SATTRLIMIT, sys-attribute parallels to
UTAGLIMIT/UATTRLIMIT.
These may be useful if only uattrs are supported, or vice-versa.
- Dropped UATTRLIMIT/SATTRLIMIT to 255 bytes.
This feels extreme, but matches NAMELIMIT. These _should_ be small,
and limiting the uattr/sattr size to a single-byte leads to really
nice packing of the utag+uattrsize in a single integer.
This can always be expanded in the future if this limit proves to be a
problem.
- Renamed MLEAFLIMIT -> MDIRLIMIT and (re?)introduced MTREELIMIT.
These may be useful to limiting the mtree when needed, though it's not
clear the exact use case quite yet.
The original name was a bit of a mouthful.
Also dropped the default crystal_size in the test/bench runners
block_size/4 -> block_size/8. I'm already noticing large amounts of
inflation when blocks are fragmented, though I am experimenting with a
rather small fragment_size right now.
Future benchmarks/experimentation is required to figure out good values
for these.
The attempt to implement in-rbyd data slicing, being lazily coalesced
during rbyd compaction, failed pretty much completely.
Slicing is a very enticing write strategy, getting both minimal overhead
post-compaction and fast random write speeds, but the idea has some
fundamental conflicts with how we play out attrs post-compaction.
This idea might work in a more powerful filesystem, but brings back the
need to simulate rbyds in RAM, which is something I really don't want to
do (complex, bug-prone, likely adds code cost, may not even be tractable).
So, third time's the charm?
---
This new write strategy writes only datas and bptrs, and avoids dagging
by completely rewriting any regions of data larger than a configurable
crystallization threshold.
This loses most of the benefits of data crystallization, random writes
will now usually need to rewrite a full block, but as a tradeoff our
data at rest is always stored with optimal overhead.
And at least data crystallization still saves space when our data isn't
block aligned, or in sparse files. From reading up on some other
filesystem designs it seems this is a desirable optimization sometimes
referred to as "tail-packing" or "block suballocation"
Some other changes from just having more time to think about the
problem:
1. Instead of scanning to figure out our current crystal size, we can
use a simple heuristic of 1. look up left block, 2. look up right
block, 3. assume any data between these blocks contribute to our
current crystal.
This is just a heuristic, so worst case you write the first and last
byte of a block which is enough to trigger compaction into a block.
But on the plus side this avoids issues with small holes preventing
blocks from being formed.
This approach brings the number of btree lookups down from
O(crystallize_size) to 2.
2. I've gone ahead and dropped the previous scheme of coalesce_size
+ fragment_size and instead adopted a single fragment_size that
controls the size of, well, fragments, i.e. data elements stored
directly in trees.
This affects both the inlined shrub as well as fragments stored in
the inner nodes of the btree. I believe it's very similar to what is
often called "pages" in logging filesystems, though I'm going to
avoid that term for now because it's a bit overloaded.
Previously, neighboring writes that, when combined, would exceed our
coalesce_size, they just weren't combined. Now they are combined up
to our fragment size, potentially splitting the right fragment.
Before (fragment_size=8):
.---+---+---+---+---+---+---+---.
| 8 bytes |
'---+---+---+---+---+---+---+---'
+
.---+---+---+---+---.
| 5 bytes |
'---+---+---+---+---'
=
.---+---+---+---+---+---+---+---+---+---.
| 5 bytes | 5 bytes |
'---+---+---+---+---+---+---+---+---+---'
After:
.---+---+---+---+---+---+---+---.
| 8 bytes |
'---+---+---+---+---+---+---+---'
+
.---+---+---+---+---.
| 5 bytes |
'---+---+---+---+---'
=
.---+---+---+---+---+---+---+---+---+---.
| 8 bytes |2 bytes|
'---+---+---+---+---+---+---+---+---+---'
This leads to better fragment alignment (much like our block
strategy), and minimizes tree overhead.
Any neighboring data to the right is only coalesced if it fits in the
current fragment, or would be rewritten (carved) anyways, to avoid
unnecessary data rewriting.
For example (fragment_size=8):
.---+---+---+---+---+---+---+---+---+---+---+---+---+---.
| 6 bytes | 6 bytes |2 bytes|
'---+---+---+---+---+---+---+---+---+---+---+---+---+---'
+
.---+---+---+---+---.
| 5 bytes |
'---+---+---+---+---'
=
.---+---+---+---+---+---+---+---+---+---+---+---+---+---.
| 8 bytes | 4 bytes |2 bytes|
'---+---+---+---+---+---+---+---+---+---+---+---+---+---'
Other than these changes this commit is mostly a bunch of carveshrub
rewriting again, which continues to be nuanced and annoying to get
bug free.
- -> lfsr_shrub_t
- -> lfsr_tree_t
The idea here is to adopt "shrub" as an umbrella term for the
shrub/sprout union, and "tree" as an umbrella term for the bptr/btree
union. I think this is a bit better than calling shrub/sprout "inlined"
which is a _very_ overloaded term in this codebase (inlined in the tree?
the mdir? inlined in the C struct?).
This is a pretty big rewrite, but is necessary to avoid "dagging".
"Dagging" (I just made this term up) is when you transform a pure tree
into a directed acyclic graph (DAG). Normally DAGs are perfectly fine in
a copy-on-write system, but in littlefs's cases, it creates havoc for
future block allocator plans, and it's interaction with parity blocks
raises some uncomfortable questions.
How does dagging happen?
Consider an innocent little btree with a single block:
.-----.
|btree|
| |
'-----'
|
v
.-----.
|abcde|
| |
'-----'
Say we wanted to write a small amount of data in the middle of our
block. Since the data is so small, the previous scheme would simply
inline the data, carving the left and right sibling (in the case the
same block) to make space:
.-----.
|btree|
| |
'-----'
.' v '.
| c' |
'. .'
v v
.-----.
|ab de|
| |
'-----'
Oh no! A DAG!
With the potential for multiple pointers to reference the same block in
our btree, some invariants break down:
- Blocks no longer have a single reference
- If you remove a reference you can no longer assume the block is free
- Knowing when a block is free requires scanning the whole btree
- This split operation effectively creates two blocks, does that mean
we need to rewrite parity blocks?
---
To avoid this whole situation, this commit adopts a new crystallization
algorithm.
Instead of allowing crystallization data to be arbitrarily fragmented,
we eagerly coalesce any data under our crystallization threshold, and if
we can't coalesce, we compact everything into a block.
Much like a Knuth heap, simply checking both siblings to coalesce has
the effect that any data will always coalesce up to the maximum size
where possible. And when checking for siblings, we can easily find the
block alignment.
This also has the effect of always rewriting blocks if we are writing a
small amount of data into a block. Unfortunately I think this is just
necessary in order to avoid dagging.
At the very least crystallization is still useful for files not quite
block aligned at the edges, and sparse files. This also avoids concerns
of random writes inflating a file via sparse crystallization.
Still needs testing, though the byte-level fuzz tests were already causing
blocks to crystallize. I noticed this because of test failures which are
fixed now.
Note the block allocator currently doesn't understand file btrees. To
get the current tests passing requires -DDISK_SIZE=16777216 or greater.
It's probably also worth noting there's a lot that's not implemented
yet! Data checksums and write validation for one. Also ecksums. And we
should probably have some sort of special handling for linear writes so
linear writes (the most common) don't end up with a bunch of extra
crystallizing writes.
Also the fact that btrees can become DAGs now is an oversight and a bit
concerning. Will that work with a closed allocator? Block parity?
Added lfsr_bptr_t to represent block pointers (maybe we should rename
mblocks back to mptr), added fetching of btrees/bptrs in
lfsr_file_opencfg, added estimate tracking to our shrubs so we actually
know when to create a btree, and implemented most of the high-level
btree logic.
It's not working yet, but the biggest idea introduced here is how we
handle block alignment.
See, we really don't want awkward btree topologies to form where small
amounts of data get stuck between blocks:
.-----.--.-----.
| | | |
| | | |
'-----'--'-----'
This is wasteful, as the middle bit of data either gets represented as a
full block with its data partially covered, or as data inlined in the
btree, which comes with ~2x overhead.
The solution here is to scan for a block on either the left or right to
derive our block alignment from.
Unfortunately, since our sibling blocks could have been carved, this
requires scanning all the way from pos-2*B+1 to pos+2*B-1, a total of
4*B-2, to make sure we find a sibling if there is one.
worst case left worst case right
.-----.-----. .-----.-----.
| xxxx| | |p |xxxxx|
|xxxxx| p| | |xxxx |
'-----'-----' '-----'-----'
'----+----' '----+----'
pos-2*bs+1 pos+2*bs-1
Fortunately, at this stage, data should have had many chances to
coalesce, so hopefully the actual scan overhead should be much smaller
in practice.
Writing data to a file linearly, for example, only needs a single lookup
to find the previous block.
Turns out it's hard to test file holes without seek.
It's interesting to note most of seek's buffer flush work actually
occurs lazily in lfsr_file_write, so lfsr_file_seek turns out to be a
relatively simple function.
- coalesce_size - The amount of data allowed to coalesce into single
data entries.
- crystallize_size - How much data is allowed to be written to btree
inner nodes before needing to be compacted into a block.
Also deduplicated the test config is something I've been wanting to do
for a while. It doesn't make sense to need to modify several different
instantiations of lfs_config every time a config option is added or
removed...
Christopher Haster