If only for consistency with DISK_GEOMETRY.
The main reason to keep BENCH_PERBYTE around is to help debug/sanity
check the more complex bus+buffer sim. For that purpose it makes sense
to be able to easily switch modes.
The only downside is if it's more difficult to introduce -DDISK_SIM=1 at
runtime vs compile-time, but eh. Consistency wins.
--no-internal has already proven useful for skipping internal tests for
refactoring, so it makes sense to add --no-reentrant/fuzz flags as well.
--no-fuzz seems particularly useful for when you want to skip the less
targeted fuzz tests:
- with fuzz tests: 634616/634616 passed, in 1239.90s
- with --no-fuzz: 85434/85434 passed, in 423.41s
I also added runtime variants to test/bench_runner and test/bench.py.
These may be useful to skip tests without needing to recompile the
runner.
---
Also tweaked -s/--step to filter permutations in any --list-* flags, for
consistency.
This has been adopted in external benchmarks for a while, as it manages
to push sequential write performance into a much better region of the
diminishing-returns curve.
But hey! Don't take my word for it, let's see the results from our new
bench_runner for the first time:
NOR throughput cs=1/8 cs=1/16
bench_wt_seq+write 15180.0 29402.6 (+93.7%)
bench_wt_random+write 876.2 957.3 (+9.3%)
bench_wt_logging+write 2001.0 2153.4 (+7.6%)
bench_wt_many+write 453.6 453.6 (+0.0%)
NAND throughput cs=1/8 cs=1/16
bench_wt_seq+write 21778.7 22330.1 (+2.5%)
bench_wt_random+write 3583.0 3637.2 (+1.5%)
bench_wt_logging+write 10855.1 10977.1 (+1.1%)
bench_wt_many+write 68.2 68.2 (+0.0%)
Though this doesn't really capture the tradeoffs related to file tails,
storage usage, etc.
In theory sequential writes are happy to start crystallizing as soon as
any data is written, but this leads to significant waste anytime you're
not going to write most of a block.
This gives us much more room for activities.
It makes sense to keep the test disk small: easier parallelization,
heavier emubd with more test features, and if you're running into space
issues in a test, that usually just means you need to be more creative
with how the test is setup.
But for benches, we're interested what happens when we throw a ton of
data at the system.
Also defaulted to noop erases. 0xff erases behave more predictably,
which is useful for testing. But for benching, less work is faster.
This shows an interesting strategy difference between the test_runner
and bench_runner.
In the test_runner we default to the least-stress configuration, to
minimize bugs unrelated to the current test. But the resulting
configuration is unrealistic, as most use cases on flash will probably
want wear-leveling.
In the bench_runner, we should use a more realistic configuration, so
setting BLOCK_RECYCLES=100 by default makes sense.
So now you can easily run multiple/specific geometries without
recompiling the bench runner:
./scripts/bench.py -DDISK_GEOMETRY=0,1
But note by default we only simulate NOR flash. Spitting out multiple
results by default is confusing.
---
Previously this was possible by either compiling multiple bench runners
(with -DBENCH_NAND), or by explicit specifying full the geometry
(-DREAD_SIZE, -DPROG_SIZE, ..., -DREAD_TIMING, ...) at runtime, but both
were clunky and annoying to parameterize.
DISK_GEOMETRY make it easy, fits well with DISK_SIZE, and adds a field
to help identify the geometry in later scripts.
I considered filling out test_defines.h with multiple geometries as
well, but decided against it. The current idea behind test_runner is to
not test specific geometries, but to instead let individual suites/cases
iterate through the specific READ_SIZEs, PROG_SIZEs, etc, that are
relevant. Still, added DISK_GEOMETRY to test_defines.h for consistency,
but it doesn't actually control anything.
After letting it sit for a bit, the previous byte+op sim comes across as
overly clever in a way that is counter-productive. This is highlighted
by erase-timing scaling in a confusing way when per-op.
Fortunately, with a bit of tweaking, we can instead model the bd sim as
separate bus+buffer timings. This seems more intuitive and is closer to
how the actual hardware works.
---
In the bus+buffer model, bd operations are simulated using two sets of
timing estimates:
buffer timings (nor) bus timings (nor)
read_timing (0) readed_timing (40 ns/B)
prog_timing (1563 ns/B) progged_timing (19 ns/B)
erase_timing (10986 ns/B) erased_timing (0)
Bus timings are a simple multiplier of the bytes read/progged/erased,
while buffer timings are rounded up + aligned to the nearest "width":
bd geometry (nor) bd buffers (nor)
read_size (1 B) read_width (1 B)
prog_size (1 B) prog_width (256 B)
erase_size (4096 B) erase_width (4096 B)
For most purposes, the width should just be the device's read/prog/erase
buffer, but I went with the name width to try to keep it generic and
avoid confusion with "buffer" elsewhere in the codebase.
Some notes:
- Like the byte+op sim, the bus+buffer sim allows penalizing small
operations without artificially limiting what operations are possible.
- Because buffer timings depend on read/prog/erase alignment, there's no
simple equation from ops+bytes to bus+buffer. But as a tradeoff, this
new sim more accurately penalizes unaligned operations.
- All timings are still kept as per-byte instead of per-width. This has
proven to be more flexible when benchmarking, as you usually what
timings to scale with the relevant operation.
- Currently this implemented by changing reads/progs/erases to track the
number of "widths" read/progged/erased after alignment. Which makes
the simtime formula roughly:
simtime = reads*read_width*read_timing + readed*readed_timing
(per-butter) (per-bus)
I considered keeping separate counters for calls (read_calls/
prog_calls/erase_calls?), but not sure there's a good reason to. The
theory behind these widths is there no functional difference between
one big call vs multiple width sized calls, though maybe they would be
useful for debugging?
We can always add these later if they turn out to be useful.
- When widths are disable (0), reads/progs/erases reverts to the number
of read/prog/erase calls.
This is the behavior when BENCH_SIMPLE is defined at compile-time.
Now that csv.py's accumulate/delta functions make it easy to switch
between delta/cumulative results, we might as well make the default
results consistent.
The previous difference between n/bench_runtime vs bench_readed/
bench_simtime risked a lot of confusion.
Note we can't use delta results for n, as it doubles as a unique index
for each probe measurement. If we want consistency the only option is
cumulative results. At least that makes the decision easy.
Having BENCH/TEST_NAND ifdefs that enable the relevant timings, but
_not_ the relevant geometry, is certainly a choice.
Defaulting to NAND geometry when BENCH/TEST_NAND is defined is more
useful, if only for minimizing confusion.
No idea how this ended up with the wrong url! I only noticed when tSE
didn't match what was in the datasheet (expected 45ms, found 50ms).
Ugh. I've been copying this url around for a while now without noticing,
so this is not the only repo that needs fixing...
Initial results with the new timing calculations looked weird. Turns
out different block sizes perform surprisingly when they all cost the
same!
Fortunately, erases are the one operation where per-byte vs per-op
timing doesn't really matter, so reverting to only per-byte timing
solves this problem. Now, erasing 2 4KiB blocks should take the same
time as 1 8KiB block, instead of twice as long.
---
Arguably, erase timing shouldn't be _strictly_ linear w.r.t. block size.
There's a reason denser storage usually ends up with larger block sizes
after all. But preventing the block size from messing with per-byte
timings is much more interesting from a filesystem design perspective.
It also matches the behavior of artificially increasing block size to
reduce block allocator pressure.
Unfortunately, this also raises concerns with read/prog timing when
varying geometry is involved... Should we stick to the per-byte timing
in such cases? Is there a better timing model out there without too much
additional complexity?
The main changes:
- Added TEST_DEFINES and BENCH_DEFINES to allow overriding the default
test/bench define header:
-DTEST_DEFINES=my_test_defines.h
Note these are VERY different from LFS_DEFINES upstream. They aren't a
typical header file, and are included multiple times with various
query macros.
It's hacky, but works surprisingly well.
Or maybe I'll just do anything to avoid having to write multiline
macros. Ugh, backslashes.
- Moved more logic into bench/test_defines.h, including everything
needed to integrate other filesystems out-of-tree.
This mostly meant moving all of the cfg initialization logic into its
own query macro (replacing the BENCH/TEST_CFG field macros).
But this also includes a bunch of small tweaks encountered while trying
to get external benchmarks running again.
The external benchmarks include several other filesystems (littlefs2,
SPIFFS, Yaffs2), and I'm hoping this injectable/queryable header thing
will do a good job at avoiding a maintenance headache. (At least a
better job than forking bench_runner.c, which was the previous
solution.)
These have been battle-tested in external benchmarks, and have proven
useful for finding a runtime estimate on stack+heap usage.
Of course, to be realistic they need to be cross-compiled and run under
QEMU (which does work!), but even on x86_64 they provide a nice insight
into RAM usage. In practice the only real difference is pointer width
anyways.
---
Enabled by default for the bench runner, these are available if
TEST/BENCH_YES_HEAP and/or TEST/BENCH_YES_STACK are defined.
(This default is provided by the Makefile. At least heap measurements
rely on linker flags, so it probably doesn't make sense to default
enable in the bench runner itself.)
Stack vs heap rely on slightly different mechanisms:
- Stack: Uses GCC's __builtin_frame_address(0) to measure the current
stack usage on entry to every bd operation.
- Heap: Relies on GCC's -Wl,--wrap flags to intercept every malloc/free
call, to track the current heap usage.
These are available via bench/test macros:
- BENCH_STACK() - Maximum stack usage of the current run
- BENCH_STACK_CURRENT() - Current stack usage
- BENCH_HEAP() - Maximum heap usage of the current run
- BENCH_HEAP_CURRENT() - Current heap usage
Note BENCH_STACK_CURRENT() can be useful for separating out the bench's
ctx from total stack usage, similarly to our static analysis.
---
One surprising outcome is that these heap hooks trivially implement a
memory leak detector. Maybe that could be useful in the test_runner as a
cheaper alternative to Valgrind?
kiwibd has been used extensively in external benchmarks, it makes sense
to make it the default bd for the bench runner:
- test_runner - defaults to emubd - more testing features
- bench_runner - defaults to kiwibd - lighter-weight disks
The benefit of kiwibd is the disk is just one big blob of RAM, so
basically no overhead. This is important when benchmarking on multi-GiB
disks.
emubd is much heavy, but as a tradeoff can do quite a bit more:
bad-block simulation, wear simulation, snapshotting, etc.
---
In theory the bd used by each runner can be controlled at compile-time
by defining -DBENCH_EMUBD, etc, but I have a feeling no one will ever
use this.
This needed a different name, and "bench probe" is sort of reminiscent
of the "debug probes" you can use to measure things in the real world.
Maybe this is just my embedded engineering background poking through,
but honestly anything is better than a single char m, especially for a
non-integer field.
This was a funny issue for external benchmarking, where we've focused
mostly on throughput benchmarking so far.
The current throughput approach is to run a benchmark for a given
simtime, and record the number of bytes written after. This is great for
allowing benchmarks to fail gracefully, but doesn't really work with the
current bench runner, which expected a known n in BENCH_START.
We can work around this by calling BENCH_START/STOP a second time
(making a mess of later scripts), but it would be nice if this was fixed
in the bench runner.
---
Humorously, BENCH_START just stores n to be printed out when BENCH_STOP
is called, so this was an easy fix.
This is based on some work in external benchmarks. What's worked well
there is emulating a global simtime based on per-byte estimates.
This moves the emulated simtime into emubd/kiwibd, and extends the idea
with both per-byte and per-op timing estimates for hopefully more
realistic results.
---
The problem is how NAND flash reads work.
Per-byte timing estimates are surprisingly accurate for NOR flash. There
is some overhead for sending the address, but it's mostly dominated by
bus cost (~20ns/B [1]).
NAND flash, on the otherhand, technically does support byte-level reads,
but first needs to read into 2KiB buffer. Surprisingly, these are pretty
close in cost (~19ns/B bus [2] vs ~12ns/B buffer [2]).
This close-ness makes modeling NAND flash difficult. If we set
read_size=1, we risk hiding the cost of small reads, which littlefs3 is
full of (rbyd lookups). If we set read_size=2048, we unfairly penalize
littlefs3 for the same reason.
---
The solution here is to expose both per-byte and per-op timing
estimates. This lets you model NAND reads using two data points:
^
| realtime --> ...............o
| : .....'''' :
| ...............:'''' ^ :
| :....''''' | :
| ..........:::::: simtime :
| .....:'''' :
|o....:::::.....: :
|: :
|: :
+:-----------------------------------------------------------:>
min read max read
Where:
bus_timing = 19ns
buffer_timing = 25us
buffer_size = 2KiB
erase_size = 128KiB
min_read = buffer_timing
max_read = (erase_size/buffer_size)*buffer_timing - buffer_timing
read_timing = min_read
readed_timing = ((max_read - min_read)/erase_size) + bus_timing
simtime = reads*read_timing + readed*readed_timing
(per-op) (per-byte)
This should correctly penalize small reads without complicating
emubd/kiwibd too much.
That's the idea anyways! It will take some use to understand if this is
a reasonable approach.
As a plus, this is a superset of the per-byte model, so both can be used
for realistic vs idealistic simulations (and to test the bus+buffer
model itself).
1: https://www.winbond.com/resource-files/W25Q256JV%20SPI%20RevQ%2002072025%20Plus.pdf
2: https://www.winbond.com/resource-files/W25N01GV%20Rev%20R%20070323.pdf
The big TEST_IMPLICIT_DEFINES and TEST_CFG macros have been a big
pain-in-the-ass to maintain. Mostly due to C preprocessor annoyances
(bleh escaped newlines) and no-ifdef workarounds, which make a real mess
of things.
This does two things:
1. Moves all the defines out of test_runner.h and into test_defines.h
(same for benches).
2. Inverts the include logic such that test_defines.h gets included many
times with various "query macros" defined.
Currently just two, but can easily add more:
1. TEST_DEFINE(name, value) - name and default value for a define
2. TEST_CFG(name, value) - name and value for a cfg field
This seems to work surprisingly well. It solves all of the above C
preprocessor issues, and provides a flexible method for defining test
defines.
Note an important part of making this work is that test_defines.h
expands to an empty string by default.
This was resulting in memory leak warnings from Valgrind, which were
getting in the way of debugging an unrelated uninitialized memory issue.
We normally wouldn't care about this sort of bounded memory leaks, but
in this case Valgrind can't tell if the memory leak is from the runner
or filesystem, errors, and prevents other tests from running. Just to be
more annoying, this only triggered when overriding defines, which is
something you do exactly when you are trying to debug something.
Fortunately, with a bit of typecasting we still have access to the
allocated value arrays (type-stripped due to opaque test_define_t), and
can clean up the relevant memory.
The original motivation for making LFS3_PREERASE opt-out, is that it
makes sense for LFS3_GBMAP to bring in all gbmap-related features
(PREERASE, BADBLOCKS (future)). However, after a bit of use, I think
this just complicates our ifdef logic too much.
So instead, LFS3_PREERASE is now opt-in, with the intention of making
all ifdefs relative only to the default build. I think this will make it
easier to reason about ifdefs, at least internally.
Eventually, I want to look into alternative default builds (LFS3_BIGGER,
LFS3_BIGGERR, ..., LFS3_BIGGEST), which would provide an alternative way
to enable all gbmap-related features. Though these builds have a
high-risk of bikeshedding (LFS3_GC?), so we'll see.
---
That being said, the main ergonomic improvement was probably adding
a #error, so we don't have to check ifdef GBMAP everywhere.
Maybe this should be extended to LFS3_RDONLY? Or maybe not, LFS3_RDONLY
is a bit of a special case.
No code changes:
code stack ctx
before: 35144 2136 660
after: 35144 (+0.0%) 2136 (+0.0%) 660 (+0.0%)
code stack ctx
gbmap before: 38380 2144 776
gbmap after: 38380 (+0.0%) 2144 (+0.0%) 776 (+0.0%)
code stack ctx
preerase before: 38920 2168 796
preerase after: 38920 (+0.0%) 2168 (+0.0%) 796 (+0.0%)
Allocating pre-erased blocks gets quite complicated due to our
restricted flash model, but at least the actual pre-erasing is
relatively straightforward:
- We keep track of known preerased state in lfs3->gbmap.preeraser.
- If LFS3_GC_PREERASE is provided during gc work, we increment the
preeraser's known window by scanning the gbmap.
- Any BMFREE ranges we find, we erase a block at a time, and store the
resulting ecksum in a BMERASED range in the gbmap.
- We keep track of how many blocks we erased, and stop early if this
exceeds cfg.gc_preerase_count. This just lets users tune how many
blocks to preerase in case something (?) prevents preerased blocks
from being used.
Some notes:
- We don't really do anything with ranges in lfs3_alloc_preerase. In
theory we could bulk in erase to minimize the number of commits to the
gbmap, but we expect erase to dominate, so this probably isn't worth
it.
And if erase doesn't dominate, why would you bother pre-erasing
blocks?
- Preerasing isn't really a traversal operation, and is managed by a
sort of secondary state machine in lfs3_fs_gc_.
This also means lfs3_trv_read with LFS3_T_PREERASE does nothing, but I
guess that is ok? It's tempting to try to make lfs3_trv_read also
preerase, but it's unclear what block it should return -- it's
probably the wrong API.
- Introducing ecksums actually went quite a bit smoother than I
expected. Though it helps ecksums are the only optional payload, no
type punning or anything.
Ecksums do muddy the gbmap's design a bit, unfortunately. The main
issue being that we can only merge BMERASED ranges with equal ecksums.
This makes BMERASED ranges less compressable than the others, and may
be one reason to limit cfg.gc_preerase_count.
However:
1. This is where I think it's useful to emphasize that the gbmap's
responsibility is to track _free_ blocks, in-use blocks are
secondary.
When allocating, we're going to stop at the first BMFREE/BMERASED,
but may need to skip over an unbounded number of BMINUSE/BMBAD
blocks. So the compressability of BMFREE/BMERASED ranges should
have less of an impact on block allocation.
2. In practice, most flash uses consistent erase values, so the
resulting ecksums will probably be compressable. The exceptions are
noop-erases (SD/eMMC, RAM, NVRAM, etc), and encryption with block
address permutation?
Though noop-erases are a pretty big exception.
Code changes:
code stack ctx
before: 35116 2136 660
after: 35116 (+0.0%) 2136 (+0.0%) 660 (+0.0%)
code stack ctx
gbmap+np before: 38040 2136 776
gbmap+np after: 38188 (+0.4%) 2144 (+0.4%) 776 (+0.0%)
code stack ctx
gbmap+yp before: 38040 2136 776
gbmap+yp after: 38608 (+1.5%) 2144 (+0.4%) 796 (+2.6%)
This extends the hopefully useful LFS3_GC_ALL flag alias to the other
functions, without trying to figure out what the heck LFS3_M_ALL should
mean semantically.
Current definitions:
// an alias for ck work
*_CK 0x00003000 ---- ---- ---- ---- --11 ---- ---- ----
// an alias for all possible gc work
*_GC 0x00003b00 ---- ---- ---- ---- ++11 11+1 ---- ----
*_MKCONSISTENT 0x00000100 ---- ---- ---- ---- ---- ---1 ---- ----
*_LOOKAHEAD 0x00000200 ---- ---- ---- ---- ---- --1- ---- ----
*_PREERASE+ 0x00000400 ---- ---- ---- ---- ---- -+-- ---- ----
*_COMPACT 0x00000800 ---- ---- ---- ---- ---- 1--- ---- ----
*_CKMETA 0x00001000 ---- ---- ---- ---- ---1 ---- ---- ----
*_CKDATA 0x00002000 ---- ---- ---- ---- --1- ---- ---- ----
*_REPAIRMETA+ 0x00004000 ---- ---- ---- ---- -+-- ---- ---- ----
*_REPAIRDATA+ 0x00008000 ---- ---- ---- ---- +--- ---- ---- ----
+ Planned
One weird artifact of this is that LFS3_GC_CK ~= LFS3_GC_CKDATA, but I'm
not sure that's a bad thing? Note this is very much not true for info
flags, LFS3_I_CK != LFS3_I_CKDATA.
No code changes.
This effectively reverts 1f824a0:
- LFS3_T_COMPACTMETA -> LFS3_T_COMPACT
- gc_compactmeta_thresh -> gc_compact_thresh
And friends.
After using LFS3_T_COMPACTMETA for a bit, I think it just adds noise
without much value. Especially when next to LFS3_T_LOOKAHEAD,
LFS3_GC_PREERASE, LFS3_M_SYNC, etc.
It's interesting that we already have some very distinct verbs for this
sort of thing based on data type (compact => metadata, garbage-collect
=> disk, compress => data).
Test/bench filters have proven to be mostly non-optional, protecting
against bad configuration that doesn't make any sense.
It's still valid to want to override test filters sometimes, but using a
more, uh, forceful verb probably makes sense here.
The shortform would conflict with -f/--fail, so no shortform flag for
this, but some argue --force should never have a shortform flag anyways.
Yeah, after using these for a bit, the RE* names were not great.
Trying LOOK* now, as an alternative that hopefully still implies the
similar behavior without needing an additional prefix for LOOKAHEAD:
- LFS3_*_RELOOKAHEAD -> LFS3_*_LOOKAHEAD
- LFS3_*_REGBMAP -> LFS3_*_LOOKGBMAP
- cfg.regbmap_thresh -> cfg.lookgbmap_thresh
- cfg.gc_relookahead_thresh -> cfg.gc_lookahead_thresh
- cfg.gc_regbmap_thresh -> cfg.gc_lookgbmap_thresh
This walks back some of the attempt at strict object namespacing in
struct lfs3_cfg:
- cfg.file_cache_size -> cfg.fcache_size
- filecfg.cache_size -> filecfg.fcache_size
- filecfg.cache_buffer -> filecfg.fcache_buffer
- cfg.gbmap_re_thresh -> cfg.regbmap_thresh
Motivation:
- cfg.regbmap_thresh now matches cfg.gc_regbmap_thresh, instead of using
awkwardly different namespacing patterns.
- Giving fcache a more unique name is useful for discussion. Having
pcache, rcache, and then file_cache was a bit awkward.
Hopefully it's also more clear that cfg.fcache_size and
filecfg.fcache_size are related.
- Config in struct lfs3_cfg is named a bit more consistently, well, if
you ignore gc_*_* options.
- Less typing.
Though this gets into pretty subjective naming territory. May revert
this if the new terms are uncomfortable after use.
So:
- cfg.gc_repoplookahead_thresh -> cfg.gc_relookahead_thresh
- cfg.gc_repopgbmap_thresh -> cfg.gc_regbmap_thresh
- cfg.gbmap_repop_thresh -> cfg.gbmap_re_thresh
- LFS3_*_REPOPLOOKAHEAD -> LFS3_*_RELOOKAHEAD
- LFS3_*_REPOPGBMAP -> LFS3_*_REGBMAP
Mainly trying to reduce the mouthful that is REPOPLOOKAHEAD and
REPOPGBMAP.
As a plus this also avoids potential confusion of "repop" as a push/pop
related operation.
To allow relaxing when LFS3_I_REPOPLOOKAHEAD and LFS3_I_REPOPGBMAP will
be set, potentially reducing gc workload after allocating only a couple
blocks.
The relevant cfg comments have quite a bit more info.
Note -1 (not the default, 0, maybe we should explicitly flip this?)
restores the previous functionality of setting these flags on the first
block allocation.
---
Also tweaked gbmap repops during gc/traversals to _not_ try to repop
unless LFS3_I_REPOPGBMAP is set. We probably should have done this from
the beginning since repopulating the gbmap writes to disk and is
potentially destructive.
Adds code, though hopefully we can claw this back with future config
rework:
code stack ctx
before: 37176 2352 684
after: 37208 (+0.1%) 2352 (+0.0%) 688 (+0.6%)
code stack ctx
gbmap before: 40024 2368 848
gbmap after: 40120 (+0.2%) 2368 (+0.0%) 856 (+0.9%)
This is an alias for all possible gc work, which is a bit more
complicated than you might think due to compile-time features (example:
LFS3_GC_REPOPGBMAP).
The intention is to make loops like the following easy to write:
struct lfs3_fsinfo fsinfo;
lfs3_fs_stat(&lfs3, &fsinfo) => 0;
lfs3_trv_t trv;
lfs3_trv_open(&lfs3, &trv, fsinfo.flags & LFS3_GC_ALL) => 0;
...
It's possible to do this by explicitly setting all gc flags, but that
requires quite a bit of knowledge from the user.
Another option is allowing -1 for gc/traversal flags, but that loses
assert protection against unknown/misplaced flags.
---
This raises more questions about the prefix naming: it feels a bit weird
to take LFS3_I_* flags, mask with LFS3_GC_* flags, and pass them as
LFS3_T_* flags, but it gets the job done.
Limiting LFS3_GC_ALL to the LFS3_GC_* namespace avoids issues with
opt-out/mode flags such as LFS3_T_RDONLY, LFS3_T_MTREEONLY, etc. For
this reason it probably doesn't make sense to add something similar to
the other namespaces.
- LFS3_T_COMPACT -> LFS3_T_COMPACTMETA
- gc_compact_thresh -> gc_compactmeta_thresh
And friends:
LFS3_M_COMPACTMETA 0x00000800 Compact metadata logs
LFS3_GC_COMPACTMETA 0x00000800 Compact metadata logs
LFS3_I_COMPACTMETA 0x00000800 Filesystem may have uncompacted metadata
LFS3_T_COMPACTMETA 0x00000800 Compact metadata logs
---
This does two things:
1. Highlights that LFS3_T_COMPACTMETA only interacts with metadata logs,
and has no effect on data blocks.
2. Better matches the verb+noun names used for other gc/traversal flags
(REPOPGBMAP, CKMETA, etc).
It is a bit more of a mouthful, but I'm not sure that's entirely a bad
thing. These are pretty low-level flags.
There's a strong argument for naming this inline_size as that's more
likely what users expect, but shrub_size is just the more correct name
and avoids confusion around having multiple names for the same thing.
It also highlights that shrubs in littlefs3 are a bit different than
inline files in littlefs2, and that this config also affects large files
with a shrubbed root.
May rerevert this in the future, but probably only if there is
significant user confusion.
And tweaked a few related comments.
I'm still on the fence with this name, I don't think it's great, but it
at least betters describes the "repopulation" operation than
"rebuilding". The important distinction is that we don't throw away
information. Bad/erased block info (future) is still carried over into
the new gbmap snapshot, and persists unless you explicitly call
rmgbmap + mkgbmap.
So, adopting gbmap_repop_thresh for now to see if it's just a habit
thing, but may adopt a different name in the future.
As a plus, gbmap_repop_thresh is two characters shorter.
I think a good rule of thumb is if you refer to some variable/config/
field with a different name in comments/writing/etc more often than not,
you should just rename the variable/config/field to match.
So yeah, gbmap_rebuild_thresh controls when the gbmap is rebuilt.
Also touched up the doc comment a bit.
Having gbmap/bmap used in different places for the same thing was
confusing. Preferring gbmap as it is consistent with other gstate (grm
queue, gcksums), even if it is a bit noisy.
It's interesting to note what didn't change:
- The BM* range tags: LFS3_TAG_BMFREE, etc. These already differs from
the GBMAP* prefix enough, and adopting GBM* would risk confusion for
actual gstate.
- The gbmap revdbg string: "bb~r". We don't have enough characters for
anything else!
- dbgbmap.py/dbgbmapsvg.py. These aren't actually related to the gbmap,
so the name difference is a good thing.
For some reason emubd had both a path argument to lfs3_emubd_create, and
a disk_path config option, with only the disk_path actually being used.
But the real curiosity is why did GCC only starting warning about it
when copied to kiwibd? path is clearly unused in lfs3_emubd_createcfg,
but no warning...
---
Anyways, not sure which one is a better API, but we definitely don't
need two APIs, so eeny meeny miny moe...
Went ahead and chose the lfs3_emubd_create path param for some
consistency with filebd.
Long story short: %zd != %jd!
This was a simple oversight when writing the bench printing code, and
easy to miss on x86_64 and other modern PCs, but the mistake becomes
very apparent when trying to bench under qemu in thumb mode!
We're not currently using these (at the moment it's unclear if the
original intention behind the treediff algorithms is worth pursuing),
and they are showing up in our heap benchmarks.
The good news is that means our heap benchmarks are working.
Also saves a bit of code/ctx in bmap mode:
code stack ctx
before: 37024 2352 684
after: 37024 (+0.0%) 2352 (+0.0%) 684 (+0.0%)
code stack ctx
bmap before: 38752 2456 812
bmap after: 38704 (-0.1%) 2456 (+0.0%) 800 (-1.5%)
At least at a proof-of-concept level, there's still a lot of cleanup
needed.
To make things work, lfs3_alloc_ckpoint now takes an mdir, which
provides the target for gbmap gstate updates.
When the bmap is close to empty (configurable via bmap_scan_thresh), we
opportunistically rebuild it during lfs3_alloc_ckpoints. The nice thing
about lfs3_alloc_ckpoint is we know the state of all in-flight blocks,
so rebuilding the bmap just requires traversing the filesystem + in-RAM
state.
We might still fall back to the lookahead buffer, but in theory a well
tuned bmap_scan_thresh can prevent this from becoming a bottleneck (at
the cost of more frequent bmap rebuilds).
---
This is also probably a good time to resume measuring code/ram costs,
though it's worth repeating the above note about the bmap work still
needing cleanup:
code stack ctx
before: 36840 2368 684
after: 36920 (+0.2%) 2368 (+0.0%) 684 (+0.0%)
Haha, no, the bmap isn't basically free, it's just an opt-in features.
With -DLFS3_YES_BMAP=1:
code stack ctx
no bmap: 36920 2368 684
yes bmap: 38552 (+4.4%) 2472 (+4.4%) 812 (+18.7%)
Note --list-suite-paths was already skipping case-less suites! I think
only -Y/--summary was an outlier.
This is consistent with test.py's matching of suite ids when no cases
are found (test_runner itself doesn't really care, it just reports no
matching cases). Though we do still compile case-less suites and include
them in the test_suites array, which may be confusing in the future.
Note this includes both the lfs3_config -> lfs3_cfg structs as well as
the LFS3_CONFIG -> LFS3_CFG include define:
- LFS3_CONFIG -> LFS3_CFG
- struct lfs3_config -> struct lfs3_cfg
- struct lfs3_file_config -> struct lfs3_file_cfg
- struct lfs3_*bd_config -> struct lfs3_*bd_cfg
- cfg -> cfg
We were already using cfg as the variable name everywhere. The fact that
these names were different was an inconsistency that should be fixed
since we're committing to an API break.
LFS3_CFG is already out-of-date from upstream, and there's plans for a
config rework, but I figured I'd go ahead and change it as well to lower
the chances it gets overlooked.
---
Note this does _not_ affect LFS3_TAG_CONFIG. Having the on-disk vs
driver-level config take slightly different names is not a bad thing.
So we now keep blocks around until they can be replaced with a single
fragment. This is simpler, cheaper, and reduces the number of commits
needed to graft (though note arbitrary range removals still keep this
unbounded).
---
So, this is a delicate tradeoff.
On one hand, not fully fragmenting blocks risks keeping around bptrs
containing very little data, depending on fragment_size.
On the other hand:
- It's expensive, and disk utilization during random _deletes_ is not
the biggest of concerns.
Note our crystallization algorithm should still clean up partial
blocks _eventually_, so this doesn't really impact random writes.
The main concerns are lfs3_file_truncate/fruncate, and in the future
collapserange/punchhole.
- Fragmenting bptrs introduces more commits, which have their own
prog/erase cost, and it's unclear how this impacts logging operations.
There's no point in fragmenting blocks at the head of a log if we're
going to fruncate them eventually.
I figure lets err on minimizing complexity/code size for now, and if
this turns out to be a mistake, we can always revert or introduce
fragmenting >1 fragment blocks as an optional feature in the future.
---
Saves a big chunk of code, stack, and even some ctx (no more
fragment_thresh):
code stack ctx
before: 37504 2448 656
after: 37024 (-1.3%) 2416 (-1.3%) 652 (-0.6%)
This one was a bit more involved.
Removes utils that are no longer useful, and made sure some of the
name/API changes over time are adopted consistently:
- lfs_npw2 -> lfs_nlog2
- lfs_tole32_ -> lfs_tole32
- lfs_fromle32_ -> lfs_fromle32
Also did another pass for lfs_ prefixes on mem/str functions. The habit
to use the naked variants of these is hard to break!
This prevents runaway O(n^2) behavior on devices with extremely large
block sizes (NAND, bs=~128KiB - ~1MiB).
The whole point of shrubs is to avoid this O(n^2) runaway when inline
files become necessarily large. Setting FRAGMENT_SIZE to a factor of the
BLOCK_SIZE humorously defeats this.
The 512 byte cutoff is somewhat arbitrary, it's the natural BLOCK_SIZE/8
FRAGMENT_SIZE on most NOR flash (bs=4096), but it's probably worth
tuning based on actual device performance.
So now crystal_thresh only controls when fragments are compacted into
blocks, while fragment_thresh controls when blocks are broken into
fragments. Setting fragment_thresh=-1 will follow crystal_thresh and
keeps the previous behavior.
These were already two separate pieces of logic, so it makes sense to
provide two separate knobs for tuning.
Setting fragment_thresh lower than crystal_thresh has some potential to
reduce hysteresis in cases where random writes push blocks close to
crystal_thresh. It will be interesting to explore this more when
benchmarking.
---
The additional config option adds a bit of code/ctx, but hopefully that
will go away in the future config rework:
code stack ctx
before: 35584 2480 636
after: 35600 (+0.0%) 2480 (+0.0%) 640 (+0.6%)
And the related config options:
- cfg->file_buffer_size -> cfg->file_cache_size
- file->cfg->buffer_size -> file->cfg->cache_size
- file->cfg->buffer -> file->cfg->cache_buffer
The original motivation to rename this to file->buffer was to better
align with what other filesystems call this, but I think this is a case
where internal consistency is more important than external consistency.
file->cache better matches lfs->pcache and lfs->rcache, and makes it
easier to read code involving both file->cache and other user-provided
buffers.
Keeping the upstream name also helps with continuity.
While I think shrub_size is probably the more correct name at a
technical level, inline_size is probably more what users expect and
doesn't require a deeper understanding of filesystem details.
The only risk is that users may think inline_size has no effect on large
files, when in fact it still controls how much of the btree root can be
inlined.
There's also the point that sticking with inline_size maintains
compatibility with both the upstream version and any future version that
has other file representations.
May revisit this, but renaming to lfs->cfg->inline_size for now.
Now that we no longer have bmoss files, inline_size and shrub_size are
effectively the same thing.
We weren't using this, so no code change, but it does save a word of
ctx:
code stack ctx
before: 36280 2576 640
after: 36280 (+0.0%) 2576 (+0.0%) 636 (-0.6%)
Christopher Haster