Wait free LW_SHARED acquisition

[Andres Freund]

Hello,

We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.

Increasing the padding to a full cacheline helps making the partitioning
of the partition space actually effective (before it's essentially
halved on a read-mostly workload), but that still leaves one with very
hot spinlocks.

So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620

master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

That's rougly 400%.

So, nice little improvement. Unless - not entirely unlikely - I fucked
up and it's fast because it's broken.

Anyway, here's the algorith I chose to implement:
The basic idea is to have a single 'uint32 lockcount' instead of the
former 'char exclusive' and 'int shared' and to use an atomic increment
to acquire the lock. That's fairly easy to do for rw-spinlocks, but a
lot harder for something like LWLocks that want to wait in the OS.

Exlusive lock acquisition:

Use an atomic compare-and-exchange on the lockcount variable swapping in
EXCLUSIVE_LOCK/1<<31/0x80000000 if and only if the current value of
lockcount is 0. If the swap was not successfull, we have to wait.


Shared lock acquisition:

Use an atomic add (lock xadd) to the lockcount variable to add 1. If the
value is bigger than EXCLUSIVE_LOCK we know that somebody actually has
an exclusive lock, and we back out by atomically decrementing by 1
again.
If so, we have to wait for the exlusive locker to release the lock.

Queueing & Wakeup:

Whenever we don't get a shared/exclusive lock we us nearly the same
queuing mechanism as we currently do. While we probably could make it
lockless as well, the queue currently is still protected by the good old
spinlock.

Relase:

Use a atomic decrement to release the lock. If the new value is zero (we
get that atomically), we know we have to release waiters.

And the real world:

Now, as you probably have noticed, naively doing the above has two
glaring race conditions:

1) too-quick-for-queueing:
We try to lock using the atomic operations and notice that we have to
wait. Unfortunately until we have finished queuing, the former locker
very well might have already finished it's work.

2) spurious failed locks:
Due to the logic of backing out of shared locks after we unconditionally
added a 1 to lockcount, we might have prevented another exclusive locker
from getting the lock:
1) Session A: LWLockAcquire(LW_EXCLUSIVE) - success
2) Session B: LWLockAcquire(LW_SHARED) - lockcount += 1
3) Session B: LWLockAcquire(LW_SHARED) - oops, bigger than EXCLUSIVE_LOCK
4) Session B: LWLockRelease()
5) Session C: LWLockAcquire(LW_EXCLUSIVE) - check if lockcount = 0, no. wait.
6) Session B: LWLockAcquire(LW_SHARED) - lockcount -= 1
7) Session B: LWLockAcquire(LW_SHARED) - wait

So now we can have both B) and C) waiting on a lock that nobody is
holding anymore. Not good.


The solution:
We use a two phased attempt at locking:
Phase 1: Try to do it atomically, if we succeed, nice
Phase 2: Add us too the waitqueue of the lock
Phase 3: Try to grab the lock again, if we succeed, remove ourselves
from the queue
Phase 4: Sleep till wakeup, goto Phase 1

This protects us against both problems from above:
1) Nobody can release too quick, before we're queued, after Phase 2 since we're already
queued.
2) If somebody spuriously got blocked from acquiring the lock, they will
get queued in Phase 2 and we can wake them up if neccessary.

Now, there are lots of tiny details to handle additionally to those, but
those seem better handled by looking at the code?
- The above algorithm only works for LWLockAcquire, not directly for
LWLockAcquireConditional where we don't want to wait. In that case we
just need to retry acquiring the lock until we're sure we didn't
disturb anybody in doing so.
- we can get removed from the queue of waiters in Phase 3, before we remove
ourselves. In that case we need to absorb the wakeup.
- Spurious locks can prevent us from recognizing a lock that's free
during release. Solve it by checking for existing waiters whenever an
exlusive lock is released.

I've done a couple of off-hand benchmarks and so far I can confirm that
everything using lots of shared locks benefits greatly and everything
else doesn't really change much. So far I've seen mostly some slight
improvements in exclusive lock heavy workloads, but those were pretty
small.
It's also very important to mention that those speedups are only there
on multi-socket machines. From what I've benchmarked so far in LW_SHARED
heavy workloads with 1 socket you get ~5-10%, 2 sockets 20-30% and
finally and nicely for 4 sockets: 350-400%.
While I did assume the difference would be bigger on 4 socket machines
than on my older 2 socket workstation (that's where the 20-30% come
from) I have to admit, I was surprised by the difference on the 4 socket
machine.

Does anybody see fundamental problems with the algorithm? The
implementation sure isn't ready for several reasons, but I don't want to
go ahead and spend lots of time on it before hearing some more voices.

So what's todo? The file header tells us:
* - revive pure-spinlock implementation
* - abstract away atomic ops, we really only need a few.
* - CAS
* - LOCK XADD
* - convert PGPROC->lwWaitLink to ilist.h slist or even dlist.
* - remove LWLockWakeup dealing with MyProc
* - overhaul the mask offsets, make SHARED/EXCLUSIVE_LOCK_MASK wider, MAX_BACKENDS

Currently only gcc is supported because I used its
__sync_fetch_and_add(), __sync_fetch_and_sub() and
__sync_val_compare_and_swap() are used. There have been reports about
__sync_fetch_and_sub() not getting properly optimized with gcc < 4.8,
perhaps we need to replace it by _and_add(-val). Given the low amount of
primitives required, it should be adaptable to most newer compilers.


Comments? Fundamental flaws? 8 socket machines?

Greetings,

Andres Freund

[Peter Geoghegan]

Post by Andres Freund
We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.

That's unfortunate. I saw someone complain about what sounds like
exactly the same issue on Twitter yesterday:

https://twitter.com/Roguelazer/status/382706273927446528

I tried to engage with him, but was unsuccessful.

[Andres Freund]

Post by Peter Geoghegan

Post by Andres Freund
We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.

That's unfortunate. I saw someone complain about what sounds like

Well, fortunately there's a solution, as presented here ;)

There's another bottleneck in the heaps of PinBuffer() calls in such
workloads, that present themselves after fixing the lwlock contention,
at least in my tests. I think I see a solution there, but let's fix this
first though ;)

Greetings,

Andres Freund

[Heikki Linnakangas]

Post by Andres Freund
Hello,
We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.
Increasing the padding to a full cacheline helps making the partitioning
of the partition space actually effective (before it's essentially
halved on a read-mostly workload), but that still leaves one with very
hot spinlocks.
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

How does that compare with simply increasing NUM_BUFFER_PARTITIONS?

- Heikki

[Andres Freund]

Hi,

Post by Heikki Linnakangas

Post by Andres Freund
We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.
Increasing the padding to a full cacheline helps making the partitioning
of the partition space actually effective (before it's essentially
halved on a read-mostly workload), but that still leaves one with very
hot spinlocks.
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

How does that compare with simply increasing NUM_BUFFER_PARTITIONS?

Heaps better. In the case causing this investigation lots of the pages
with hot spinlocks were the simply the same ones over and over again,
partitioning the lockspace won't help much there.
That's not exactly an uncommon scenario since often enough there's a
small amount of data hit very frequently and lots more that is accessed
only infrequently. E.g. recently inserted data and such tends to be very hot.

I can run a test on the 4 socket machine if it's unused, but on my 2
socket workstation the benefits of at least our simulation of the
original workloads the improvements were marginal after increasing the
padding to a full cacheline.

Greetings,

Andres Freund

[Andres Freund]

Post by Andres Freund

Post by Heikki Linnakangas

Post by Andres Freund
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

How does that compare with simply increasing NUM_BUFFER_PARTITIONS?

Heaps better. In the case causing this investigation lots of the pages
with hot spinlocks were the simply the same ones over and over again,
partitioning the lockspace won't help much there.
That's not exactly an uncommon scenario since often enough there's a
small amount of data hit very frequently and lots more that is accessed
only infrequently. E.g. recently inserted data and such tends to be very hot.
I can run a test on the 4 socket machine if it's unused, but on my 2
socket workstation the benefits of at least our simulation of the
original workloads the improvements were marginal after increasing the
padding to a full cacheline.

Ok, was free:

padding + 16 partitions:
tps = 147884.648416

padding + 32 partitions:
tps = 141777.841125

padding + 64 partitions:
tps = 141561.539790

padding + 16 partitions + new lwlocks
tps = 601895.580903 (yeha, still reproduces after some sleep!)


Now, the other numbers were best-of-three, these aren't, but I think
it's pretty clear that you're not going to see the same benefits. I am
not surprised...
The current implementation of lwlocks will frequently block others, both
during acquiration and release of locks. What's even worse, trying to
fruitlessly acquire a spinlock will often prevent releasing it because
we need the spinlock during release.
With the proposed algorithm, even if we need the spinlock during release
of an lwlock because there are queued PGPROCs, we will acquire that
spinlock only after already having released the lock...

Greetings,

Andres Freund

[Andres Freund]

Post by Andres Freund

Post by Andres Freund

Post by Heikki Linnakangas

Post by Andres Freund
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

How does that compare with simply increasing NUM_BUFFER_PARTITIONS?

Heaps better. In the case causing this investigation lots of the pages
with hot spinlocks were the simply the same ones over and over again,
partitioning the lockspace won't help much there.
That's not exactly an uncommon scenario since often enough there's a
small amount of data hit very frequently and lots more that is accessed
only infrequently. E.g. recently inserted data and such tends to be very hot.
I can run a test on the 4 socket machine if it's unused, but on my 2
socket workstation the benefits of at least our simulation of the
original workloads the improvements were marginal after increasing the
padding to a full cacheline.

tps = 147884.648416
tps = 141777.841125
tps = 141561.539790
padding + 16 partitions + new lwlocks
tps = 601895.580903 (yeha, still reproduces after some sleep!)

Pgbench numbers for writes on the machine (fsync = off,
synchronous_commit = off):
padding + 16 partitions:
tps = 8903.532163
vs
padding + 16 partitions + new lwlocks
tps = 13324.080491

So, on bigger machines the advantages seem to be there for writes as
well...

Greetings,

Andres Freund

[Bernd Helmle]

--On 27. September 2013 09:57:07 +0200 Andres Freund

Post by Andres Freund
tps = 147884.648416
tps = 141777.841125
tps = 141561.539790
padding + 16 partitions + new lwlocks
tps = 601895.580903 (yeha, still reproduces after some sleep!)

Hmm, out of interest and since i have access to a (atm) free DL580 G7 (4x
E7-4800 10core) i've tried your bench against this machine and got this
(best of three):

HEAD (default):

tps = 181738.607247 (including connections establishing)
tps = 182665.993063 (excluding connections establishing)

HEAD (padding + 16 partitions + your lwlocks patch applied):

tps = 269328.259833 (including connections establishing)
tps = 270685.666091 (excluding connections establishing)

So, still an improvement but far away from what you got. Do you have some
other tweaks in your setup?
--
Thanks

Bernd

[Andres Freund]

Hi,

Post by Bernd Helmle
tps = 181738.607247 (including connections establishing)
tps = 182665.993063 (excluding connections establishing)
tps = 269328.259833 (including connections establishing)
tps = 270685.666091 (excluding connections establishing)
So, still an improvement but far away from what you got. Do you have some
other tweaks in your setup?

The only relevant setting changed was -c shared_buffers=1GB, no other
patches applied. At which scale did you pgbench -i?

Your processors are a different microarchitecture, I guess that could
also explain some of the difference. Any chance you could run a perf record -ag
(after compiling with -fno-omit-frame-pointer)?

Greetings,

Andres Freund

[Bernd Helmle]

--On 30. September 2013 19:00:06 +0200 Andres Freund

Post by Andres Freund

Post by Bernd Helmle
tps = 181738.607247 (including connections establishing)
tps = 182665.993063 (excluding connections establishing)
tps = 269328.259833 (including connections establishing)
tps = 270685.666091 (excluding connections establishing)
So, still an improvement but far away from what you got. Do you have some
other tweaks in your setup?

The only relevant setting changed was -c shared_buffers=1GB, no other
patches applied. At which scale did you pgbench -i?

I've used a scale factor of 10 (i recall you've mentioned using the same
upthread...).

Okay, i've used 2GB shared buffers, repeating with your setting i get a far
more noticable speedup:

tps = 346292.008580 (including connections establishing)
tps = 347997.073595 (excluding connections establishing)

Here's the perf output:

+ 4.34% 207112 postgres postgres [.]
AllocSetAlloc
+ 4.07% 194476 postgres libc-2.13.so [.] 0x127b33
+ 2.59% 123471 postgres postgres [.]
SearchCatCache
+ 2.49% 118974 pgbench libc-2.13.so [.] 0x11aaef
+ 2.48% 118263 postgres postgres [.]
GetSnapshotData
+ 2.46% 117646 postgres postgres [.]
base_yyparse
+ 2.02% 96546 postgres postgres [.]
MemoryContextAllocZeroAligned
+ 1.58% 75326 postgres postgres [.]
AllocSetFreeIndex
+ 1.23% 58587 postgres postgres [.]
hash_search_with_hash_value
+ 1.01% 48391 postgres postgres [.] palloc
+ 0.93% 44258 postgres postgres [.]
LWLockAttemptLock
+ 0.91% 43575 pgbench libc-2.13.so [.] free
+ 0.89% 42484 postgres postgres [.]
nocachegetattr
+ 0.89% 42378 postgres postgres [.] core_yylex
+ 0.88% 42001 postgres postgres [.]
_bt_compare
+ 0.84% 39997 postgres postgres [.]
expression_tree_walker
+ 0.76% 36533 postgres postgres [.]
ScanKeywordLookup
+ 0.74% 35515 pgbench libc-2.13.so [.] malloc
+ 0.64% 30715 postgres postgres [.]
LWLockRelease
+ 0.56% 26779 postgres postgres [.]
fmgr_isbuiltin
+ 0.54% 25681 pgbench [kernel.kallsyms] [k] _spin_lock
+ 0.48% 22836 postgres postgres [.] new_list
+ 0.48% 22700 postgres postgres [.] hash_any
+ 0.47% 22378 postgres postgres [.]
FunctionCall2Coll
+ 0.46% 22095 postgres postgres [.] pfree
+ 0.44% 20929 postgres postgres [.] palloc0
+ 0.43% 20592 postgres postgres [.]
AllocSetFree
+ 0.40% 19495 postgres [unknown] [.] 0x81cf2f
+ 0.40% 19247 postgres postgres [.]
hash_uint32
+ 0.38% 18227 postgres postgres [.] PinBuffer
+ 0.38% 18022 pgbench [kernel.kallsyms] [k] do_select
--
Thanks

Bernd

[Merlin Moncure]

Post by Bernd Helmle
--On 30. September 2013 19:00:06 +0200 Andres Freund

Post by Andres Freund

Post by Bernd Helmle
tps = 181738.607247 (including connections establishing)
tps = 182665.993063 (excluding connections establishing)
tps = 269328.259833 (including connections establishing)
tps = 270685.666091 (excluding connections establishing)
So, still an improvement but far away from what you got. Do you have some
other tweaks in your setup?

The only relevant setting changed was -c shared_buffers=1GB, no other
patches applied. At which scale did you pgbench -i?

I've used a scale factor of 10 (i recall you've mentioned using the same
upthread...).
Okay, i've used 2GB shared buffers, repeating with your setting i get a far

If Andres's patch passes muster it may end up causing us to
re-evaluate practices for the shared buffer setting. I was trying to
optimize buffer locking in the clock sweep using a different approach
and gave up after not being able to find useful test cases to
demonstrate an improvement. The main reason for this is that clock
sweep issues are masked by contention in the buffer mapping lwlocks
(as you guys noted). I *do* think clock sweep contention comes out in
some production workloads but so far have been elusive to produce in
synthetic testing. Ditto buffer pin contention (this has been
documented).

So I'm very excited about this patch. Right now in servers I
configure (even some very large ones) I set shared buffers to max 2gb
for various reasons. Something tells me that's about to change.

merlin

[Andres Freund]

Post by Bernd Helmle
--On 30. September 2013 19:00:06 +0200 Andres Freund

Post by Andres Freund

Post by Bernd Helmle
tps = 181738.607247 (including connections establishing)
tps = 182665.993063 (excluding connections establishing)
tps = 269328.259833 (including connections establishing)
tps = 270685.666091 (excluding connections establishing)
So, still an improvement but far away from what you got. Do you have some
other tweaks in your setup?

The only relevant setting changed was -c shared_buffers=1GB, no other
patches applied. At which scale did you pgbench -i?

I've used a scale factor of 10 (i recall you've mentioned using the same
upthread...).
Okay, i've used 2GB shared buffers, repeating with your setting i get a far
tps = 346292.008580 (including connections establishing)
tps = 347997.073595 (excluding connections establishing)

Could you send hierarchical profiles of both 1 and 2GB? It's curious
that the difference is that big... Even though they will be a bit big,
it'd be helpful if you pasted the output of "perf report --stdio", to
include the callers...

Greetings,

Andres Freund

[Heikki Linnakangas]

Post by Andres Freund
Hi,

Post by Heikki Linnakangas

Post by Andres Freund
We have had several customers running postgres on bigger machines report
problems on busy systems. Most recently one where a fully cached
workload completely stalled in s_lock()s due to the *shared* lwlock
acquisition in BufferAlloc() around the buffer partition lock.
Increasing the padding to a full cacheline helps making the partitioning
of the partition space actually effective (before it's essentially
halved on a read-mostly workload), but that still leaves one with very
hot spinlocks.
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065

How does that compare with simply increasing NUM_BUFFER_PARTITIONS?

Heaps better. In the case causing this investigation lots of the pages
with hot spinlocks were the simply the same ones over and over again,
partitioning the lockspace won't help much there.
That's not exactly an uncommon scenario since often enough there's a
small amount of data hit very frequently and lots more that is accessed
only infrequently. E.g. recently inserted data and such tends to be very hot.

I see. So if only a few buffers are really hot, I'm assuming the problem
isn't just the buffer partition lock, but also the spinlock on the
buffer header, and the buffer content lwlock. Yeah, improving LWLocks
would be a nice wholesale solution to that. I don't see any fundamental
flaw in your algorithm. Nevertheless, I'm going to throw in a couple of
other ideas:

* Keep a small 4-5 entry cache of buffer lookups in each backend of most
recently accessed buffers. Before searching for a buffer in the
SharedBufHash, check the local cache.

* To pin a buffer, use an atomic fetch-and-add instruction to increase
the refcount. PinBuffer() also increases usage_count, but you could do
that without holding a lock; it doesn't need to be accurate.


One problem with your patch is going to be to make it also work without
the CAS and fetch-and-add instructions. Those are probably present in
all the architectures we support, but it'll take some effort to get the
architecture-specific code done. Until it's all done, it would be good
to be able to fall back to plain spinlocks, which we already have. Also,
when someone ports PostgreSQL to a new architecture in the future, it
would be helpful if you wouldn't need to write all the
architecture-specific code immediately to get it to compile.

Did you benchmark your patch against the compare-and-swap patch I posted
earlier?
(http://www.postgresql.org/message-id/***@vmware.com). Just
on a theoretical level, I would assume your patch to scale better since
it uses fetch-and-add instead of compare-and-swap for acquiring a shared
lock. But in practice it might be a wash.

- Heikki

[Andres Freund]

Post by Heikki Linnakangas

Post by Andres Freund
Heaps better. In the case causing this investigation lots of the pages
with hot spinlocks were the simply the same ones over and over again,
partitioning the lockspace won't help much there.
That's not exactly an uncommon scenario since often enough there's a
small amount of data hit very frequently and lots more that is accessed
only infrequently. E.g. recently inserted data and such tends to be very hot.

I see. So if only a few buffers are really hot, I'm assuming the problem
isn't just the buffer partition lock, but also the spinlock on the buffer
header, and the buffer content lwlock. Yeah, improving LWLocks would be a
nice wholesale solution to that. I don't see any fundamental flaw in your
* Keep a small 4-5 entry cache of buffer lookups in each backend of most
recently accessed buffers. Before searching for a buffer in the
SharedBufHash, check the local cache.

I thought about that as well, but you'd either need to revalidate after
pinning the buffers, or keep them pinned.
I had a very hacky implementation of that, but it just make the buffer
content locks the top profile spots. Similar issue there.

It might be worthwile to do this nonetheless - lock xadd; certainly
isn't cheap, even due it's cheaper than a full spinnlock - but it's not
trivial.

Post by Heikki Linnakangas
* To pin a buffer, use an atomic fetch-and-add instruction to increase the
refcount. PinBuffer() also increases usage_count, but you could do that
without holding a lock; it doesn't need to be accurate.

Yes, Pin/UnpinBuffer() are the primary contention points after this
patch. I want to tackle them, but that seems like a separate thing to
do.

I think we should be able to get rid of most or even all LockBufHdr()
calls by
a) introducing PinBufferInternal() which increase pins but not
usage count using an atomic increment. That can replace locking headers
in many cases.
b) make PinBuffer() increment pin and usagecount using a single 64bit
atomic add if available and fit flags in there as well. Then back off
the usagecount if it's too high or even wraps around, that doesn't hurt
much, we're pinned in that moment.

Post by Heikki Linnakangas
One problem with your patch is going to be to make it also work without the
CAS and fetch-and-add instructions. Those are probably present in all the
architectures we support, but it'll take some effort to get the
architecture-specific code done. Until it's all done, it would be good to be
able to fall back to plain spinlocks, which we already have. Also, when
someone ports PostgreSQL to a new architecture in the future, it would be
helpful if you wouldn't need to write all the architecture-specific code
immediately to get it to compile.

I think most recent compilers have intrinsics for stuff for operations
like that. I quite like the API provided by gcc for this kind of stuff,
I think we should model an internal wrapper API similarly. I don't see
any new platforming coming that has a compiler without intrinsics?

But yes, you're right, I think we need a spinlock based fallback for
now. Even if it's just because nobody of us can verify the
implementations on the more obsolete platforms we claim to support.I
just didn't see it as a priority in the PoC.

Post by Heikki Linnakangas
Did you benchmark your patch against the compare-and-swap patch I posted
Just on a theoretical level, I would assume your patch to scale better since
it uses fetch-and-add instead of compare-and-swap for acquiring a shared
lock. But in practice it might be a wash.

I've tried compare-and-swap for shared acquisition and it performed
worse, didn't try your patch though as you seemed to have concluded it's
a wash after doing the unlocked test.

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Florian Pflug]

Post by Andres Freund
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065
That's rougly 400%.

Interesting. I played with pretty much the same idea two years or so ago.
At the time, I compared a few different LWLock implementations. Those
were AFAIR

A) Vanilla LWLocks
B) A + an atomic-increment fast path, very similar to your proposal
C) B but with a partitioned atomic-increment counter to further
reduce cache-line contention
D) A with the spinlock-based queue replaced by a lockless queue

At the time, the improvements seemed to be negligible - they looked great
in synthetic benchmarks of just the locking code, but didn't translate
to improved TPS numbers. Though I think the only version that ever got
tested on more than a handful of cores was C…

My (rather hacked together) benchmarking code can be found here: https://github.com/fgp/lockbench.
The different LWLock implementations live in the various pg_lwlock_* subfolders.

Here's a pointer into the relevant thread: http://www.postgresql.org/message-id/651002C1-2EC1-4731-9B29-***@phlo.org

best regards,
Florian Pflug

[Andres Freund]

Post by Florian Pflug

Post by Andres Freund
So the goal is to have LWLockAcquire(LW_SHARED) never block unless
somebody else holds an exclusive lock. To produce enough appetite for
reading the rest of the long mail, here's some numbers on a
pgbench -j 90 -c 90 -T 60 -S (-i -s 10) on a 4xE5-4620
master + padding: tps = 146904.451764
master + padding + lwlock: tps = 590445.927065
That's rougly 400%.

Interesting. I played with pretty much the same idea two years or so ago.
At the time, I compared a few different LWLock implementations. Those
were AFAIR
A) Vanilla LWLocks
B) A + an atomic-increment fast path, very similar to your proposal
C) B but with a partitioned atomic-increment counter to further
reduce cache-line contention
D) A with the spinlock-based queue replaced by a lockless queue
At the time, the improvements seemed to be negligible - they looked great
in synthetic benchmarks of just the locking code, but didn't translate
to improved TPS numbers. Though I think the only version that ever got
tested on more than a handful of cores was C…

I think you really need multi-socket systems to see the big benefits
from this. My laptop barely shows any improvements, while my older 2
socket workstation already shows some in workloads that have more
contention than pgbench -S.

From a quick look, you didn't have any sleeping queueing in at least one
of the variants in there? In my tests, that was tremendously important
to improve scaling if there was any contention. Which is not surprising
in the end, because otherwise you essentially have rw-spinlocks which
really aren't suitable for many of the lwlocks we use.

Getting the queueing semantics, including releaseOK, right was what took
me a good amount of time, the atomic ops part was pretty quick...

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Andres Freund]

Hi,

Post by Andres Freund
* - revive pure-spinlock implementation
* - abstract away atomic ops, we really only need a few.
* - CAS
* - LOCK XADD
* - convert PGPROC->lwWaitLink to ilist.h slist or even dlist.
* - remove LWLockWakeup dealing with MyProc
* - overhaul the mask offsets, make SHARED/EXCLUSIVE_LOCK_MASK wider, MAX_BACKENDS

So, here's the next version of this patchset:
1) I've added an abstracted atomic ops implementation. Needs a fair
amount of work, also submitted as a separate CF entry. (Patch 1 & 2)
2) I've converted PGPROC->lwWaiting into a dlist. That makes a fair bit
of code easier to read and reduces the size of the patchset. Also
fixes a bug in the xlog-scalability code. (Patch 3)
3) Alvaro and I updated the comments in lwlock.c. (Patch 4)

I think 2) should be committable pretty soon. It's imo a pretty clear
win in readability. 1) will need a good bit of more work.

With regard to the scalable lwlock work, what's most needed now is a
good amount of testing.

Please note that you need to 'autoreconf' after applying the patchset. I
don't have a compatible autoconf version on this computer causing the
diff to be humongous if I include those changes.

Greetings,

Andres Freund

[Andres Freund]

Post by Andres Freund
Hi,

Post by Andres Freund
* - revive pure-spinlock implementation
* - abstract away atomic ops, we really only need a few.
* - CAS
* - LOCK XADD
* - convert PGPROC->lwWaitLink to ilist.h slist or even dlist.
* - remove LWLockWakeup dealing with MyProc
* - overhaul the mask offsets, make SHARED/EXCLUSIVE_LOCK_MASK wider, MAX_BACKENDS

1) I've added an abstracted atomic ops implementation. Needs a fair
amount of work, also submitted as a separate CF entry. (Patch 1 & 2)
2) I've converted PGPROC->lwWaiting into a dlist. That makes a fair bit
of code easier to read and reduces the size of the patchset. Also
fixes a bug in the xlog-scalability code. (Patch 3)
3) Alvaro and I updated the comments in lwlock.c. (Patch 4)
I think 2) should be committable pretty soon. It's imo a pretty clear
win in readability. 1) will need a good bit of more work.
With regard to the scalable lwlock work, what's most needed now is a
good amount of testing.
Please note that you need to 'autoreconf' after applying the patchset. I
don't have a compatible autoconf version on this computer causing the
diff to be humongous if I include those changes.

Please also note that due to the current state of the atomics
implementation this likely will only work on a somewhat recent gcc and
that the performance might be slightly worse than in the previous
version because the atomic add is implemented using the CAS fallback.

Greetings,

Andres Freund

[Peter Eisentraut]

The 0002 patch contains non-ASCII, non-UTF8 characters:

0002-Very-basic-atomic-ops-implementation.patch: line 609, char 1, byte offset 43: invalid UTF-8 code

Please change that to ASCII, or UTF-8 if necessary.

[Peter Eisentraut]

This patch didn't make it out of the 2013-11 commit fest. You should
move it to the next commit fest (probably with an updated patch)
before January 15th.

[Peter Geoghegan]

Post by Andres Freund
1) I've added an abstracted atomic ops implementation. Needs a fair
amount of work, also submitted as a separate CF entry. (Patch 1 & 2)

Commit 220b34331f77effdb46798ddd7cca0cffc1b2858 caused bitrot when
applying 0002-Very-basic-atomic-ops-implementation.patch. Please
rebase.

[Andres Freund]

Hi,

Post by Peter Geoghegan

Post by Andres Freund
1) I've added an abstracted atomic ops implementation. Needs a fair
amount of work, also submitted as a separate CF entry. (Patch 1 & 2)

Commit 220b34331f77effdb46798ddd7cca0cffc1b2858 caused bitrot when
applying 0002-Very-basic-atomic-ops-implementation.patch. Please
rebase.

I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.

I plan to split the atomics patch into smaller chunks before
reposting. Imo the "Convert the PGPROC->lwWaitLink list into a dlist
instead of open coding it." is worth being applied independently from
the rest of the series, it simplies code and it fixes a bug...

Greetings,

Andres Freund

[Peter Geoghegan]

Post by Andres Freund
I plan to split the atomics patch into smaller chunks before
reposting. Imo the "Convert the PGPROC->lwWaitLink list into a dlist
instead of open coding it." is worth being applied independently from
the rest of the series, it simplies code and it fixes a bug...

For things that require a format-patch series, I personally find it
easier to work off a feature branch on a remote under the control of
the patch author. The only reason that I don't do so myself is that I
know that that isn't everyone's preference.

I have access to a large server for the purposes of benchmarking this.
On the plus side, this is a box very much capable of exercising these
bottlenecks: a 48 core AMD system, with 256GB of RAM. However, I had
to instruct someone else on how to conduct the benchmark, since I
didn't have SSH access, and the OS and toolchain were antiquated,
particularly for this kind of thing. This is Linux kernel
2.6.18-274.3.1.el5 (RHEL5.7). The GCC version that Postgres was built
with was 4.1.2. This is not what I'd hoped for; obviously I would have
preferred to be able to act on your warning: "Please also note that
due to the current state of the atomics implementation this likely
will only work on a somewhat recent gcc and that the performance might
be slightly worse than in the previous version because the atomic add
is implemented using the CAS fallback". Even still, it might be
marginally useful to get a sense of that cost.

I'm looking at alternative options, because this is not terribly
helpful. With those big caveats in mind, consider the results of the
benchmark, which show the patch performing somewhat worse than the
master baseline at higher client counts:

http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/rwlock-contention/

This is exactly what you said would happen, but at least now we have a
sense of that cost.

[Peter Geoghegan]

I thought I'd try out what I was in an immediate position to do
without having access to dedicated multi-socket hardware: A benchmark
on AWS. This was a "c3.8xlarge" instance, which are reportedly backed
by Intel Xeon E5-2680 processors. Since the Intel ARK website reports
that these CPUs have 16 "threads" (8 cores + hyperthreading), and
AWS's marketing material indicates that this instance type has 32
"vCPUs", I inferred that the underlying hardware had 2 sockets.
However, reportedly that wasn't the case when procfs was consulted, no
doubt due to Xen Hypervisor voodoo:

***@ip-10-67-128-2:~$ lscpu
Architecture: x86_64
CPU op-mode(s): 32-bit, 64-bit
Byte Order: Little Endian
CPU(s): 32
On-line CPU(s) list: 0-31
Thread(s) per core: 32
Core(s) per socket: 1
Socket(s): 1
NUMA node(s): 1
Vendor ID: GenuineIntel
CPU family: 6
Model: 62
Stepping: 4
CPU MHz: 2800.074
BogoMIPS: 5600.14
Hypervisor vendor: Xen
Virtualization type: full
L1d cache: 32K
L1i cache: 32K
L2 cache: 256K
L3 cache: 25600K
NUMA node0 CPU(s): 0-31

I ran the benchmark on Ubuntu 13.10 server, because that seemed to be
the only prominent "enterprise" x86_64 AMI (operating system image)
that came with GCC 4.8 as part its standard toolchain. This exact
setup is benchmarked here:

http://www.phoronix.com/scan.php?page=article&item=amazon_ec2_c3&num=1

(Incidentally, some of the other benchmarks on that site use pgbench
to benchmark the Linux kernel, filesystems, disks and so on. e.g.:
http://www.phoronix.com/scan.php?page=news_item&px=NzI0NQ).

I was hesitant to benchmark using a virtualized system. There is a lot
of contradictory information about the overhead and/or noise added,
which may vary from one workload or hypervisor to the next. But, needs
must when the devil drives, and all that. Besides, this kind of setup
is very commercially relevant these days, so it doesn't seem
unreasonable to see how things work out on an AWS instance that
generally performs well for this workload. Of course, I still want to
replicate the big improvement you reported for multi-socket systems,
but you might have to wait a while for that, unless some kindly
benefactor that has a 4 socket server lying around would like to help
me out.

Results:

http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/c38xlarge-rwlocks/

You can drill down and find the postgresql.conf settings from the
report. There appears to be a modest improvement in transaction
throughput. It's not as large as the improvement you reported for your
2 socket workstation, but it's there, just about.

[Andres Freund]

Post by Peter Geoghegan

Post by Andres Freund
I plan to split the atomics patch into smaller chunks before
reposting. Imo the "Convert the PGPROC->lwWaitLink list into a dlist
instead of open coding it." is worth being applied independently from
the rest of the series, it simplies code and it fixes a bug...

For things that require a format-patch series, I personally find it
easier to work off a feature branch on a remote under the control of
the patch author. The only reason that I don't do so myself is that I
know that that isn't everyone's preference.

I do to, that's why I have a git branch for all but the most trivial
branches.

Post by Peter Geoghegan
I have access to a large server for the purposes of benchmarking this.
On the plus side, this is a box very much capable of exercising these
bottlenecks: a 48 core AMD system, with 256GB of RAM. However, I had
to instruct someone else on how to conduct the benchmark, since I
didn't have SSH access, and the OS and toolchain were antiquated,
particularly for this kind of thing. This is Linux kernel
2.6.18-274.3.1.el5 (RHEL5.7). The GCC version that Postgres was built
with was 4.1.2. This is not what I'd hoped for; obviously I would have
preferred to be able to act on your warning: "Please also note that
due to the current state of the atomics implementation this likely
will only work on a somewhat recent gcc and that the performance might
be slightly worse than in the previous version because the atomic add
is implemented using the CAS fallback". Even still, it might be
marginally useful to get a sense of that cost.

I *think* it should actually work on gcc 4.1, I've since added the
fallbacks I hadn't back when I wrote the above. I've added exactly those
atomics that are needed for the scalable lwlock things (xadd, xsub (yes,
that's essentially the same) and cmpxchg).

Post by Peter Geoghegan
I'm looking at alternative options, because this is not terribly
helpful. With those big caveats in mind, consider the results of the
benchmark, which show the patch performing somewhat worse than the

I think that's actually something else. I'd tried to make some
definitions simpler, and that has, at least for the machine I have
occasional access to, pessimized things. I can't always run the tests
there, so I hadn't noticed before the repost.
I've pushed a preliminary relief to the git repository, any chance you
could retry?

Greetings,

Andres Freund

[Peter Geoghegan]

Post by Andres Freund

Post by Peter Geoghegan
I'm looking at alternative options, because this is not terribly
helpful. With those big caveats in mind, consider the results of the
benchmark, which show the patch performing somewhat worse than the

I think that's actually something else. I'd tried to make some
definitions simpler, and that has, at least for the machine I have
occasional access to, pessimized things. I can't always run the tests
there, so I hadn't noticed before the repost.

I should have been clearer on one point: The pre-rebased patch (actual
patch series) [1] was applied on top of a commit from around the same
period, in order to work around the bit rot. However, I tested the
most recent revision from your git remote on the AWS instance.

[1] http://www.postgresql.org/message-id/***@awork2.anarazel.de

[Andres Freund]

Post by Peter Geoghegan

Post by Andres Freund

Post by Peter Geoghegan
I'm looking at alternative options, because this is not terribly
helpful. With those big caveats in mind, consider the results of the
benchmark, which show the patch performing somewhat worse than the

I think that's actually something else. I'd tried to make some
definitions simpler, and that has, at least for the machine I have
occasional access to, pessimized things. I can't always run the tests
there, so I hadn't noticed before the repost.

I should have been clearer on one point: The pre-rebased patch (actual
patch series) [1] was applied on top of a commit from around the same
period, in order to work around the bit rot.

Ah. Then I indeed wouldn't expect improvements.

Post by Peter Geoghegan
However, I tested the
most recent revision from your git remote on the AWS instance.

But that was before my fix, right. Except you managed to timetravel :)

Greetings,

Andres Freund

[Peter Geoghegan]

Post by Andres Freund

Post by Peter Geoghegan
However, I tested the
most recent revision from your git remote on the AWS instance.

But that was before my fix, right. Except you managed to timetravel :)

Heh, okay. So Nathan Boley has generously made available a machine
with 4 AMD Opteron 6272s. I've performed the same benchmark on that
server. However, I thought it might be interesting to circle back and
get some additional numbers for the AWS instance already tested - I'd
like to see what it looks like after your recent tweaks to fix the
regression. The single client performance of that instance seems to be
markedly better than that of Nathan's server.

Tip: AWS command line tools + S3 are a great way to easily publish
bulky pgbench-tools results, once you figure out how to correctly set
your S3 bucket's security manifest to allow public http. It has
similar advantages to rsync, and just works with the minimum of fuss.

Anyway, I don't think that the new, third c3.8xlarge-rwlocks testset
tells us much of anything:
http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/c38xlarge-rwlocks/

Here are the results of a benchmark on Nathan Boley's 64-core, 4
socket server: http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/

Perhaps I should have gone past 64 clients, because in the document
"Lock Scaling Analysis on Intel Xeon Processors" [1], Intel write:

"This implies that with oversubscription (more threads running than
available logical CPUs), the performance of spinlocks can depend
heavily on the exact OS scheduler behavior, and may change drastically
with operating system or VM updates."

I haven't bothered with a higher client counts though, because Andres
noted it's the same with 90 clients on this AMD system. Andres: Do you
see big problems when # clients < # logical cores on the affected
Intel systems?

There is only a marginal improvement in performance on this big 4
socket system. Andres informs me privately that he has reproduced the
problem on multiple new 4-socket Intel servers, so it seems reasonable
to suppose that more or less an Intel thing.

The Intel document [1] further notes:

"As the number of threads polling the status of a lock address
increases, the time it takes to process those polling requests will
increase. Initially, the latency to transfer data across socket
boundaries will always be an order of magnitude longer than the
on-chip cache-to-cache transfer latencies. Such cross-socket
transfers, if they are not effectively minimized by software, will
negatively impact the performance of any lock algorithm that depends
on them."

So, I think it's fair to say, given what we now know from Andres'
numbers and the numbers I got from Nathan's server, that this patch is
closer to being something that addresses a particularly unfortunate
pathology on many-socket Intel system than it is to being a general
performance optimization. Based on the above quoted passage, it isn't
unreasonable to suppose other vendors or architectures could be
affected, but that isn't in evidence. While I welcome the use of
atomic operations in the context of LW_SHARED acquisition as general
progress, I think that to the outside world my proposed messaging is
more useful. It's not quite a bug-fix, but if you're using a
many-socket Intel server, you're *definitely* going to want to use a
PostgreSQL version that is unaffected. You may well not want to take
on the burden of waiting for 9.4, or waiting for it to fully
stabilize.

I note that Andres has a feature branch of this backported to Postgres
9.2, no doubt because of a request from a 2ndQuadrant customer. I have
to wonder if we should think about making this available with a
configure switch in one or more back branches. I think that the
complete back-porting of the fsync request queue issue's fix in commit
758728 could be considered a precedent - that too was a fix for a
really bad worst-case that was encountered fairly infrequently in the
wild. It's sort of horrifying to have red-hot spinlocks in production,
so that seems like the kind of thing we should make an effort to
address for those running multi-socket systems. Of those running
Postgres on new multi-socket systems, the reality is that the majority
are running on Intel hardware. Unfortunately, everyone knows that
Intel will soon be the only game in town when it comes to high-end
x86_64 servers, which contributes to my feeling that we need to target
back branches. We should do something about the possible regression
with older compilers using the fallback first, though.

It would be useful to know more about the nature of the problem that
made such an appreciable difference in Andres' original post. Again,
through private e-mail, I saw perf profiles from affected servers and
an unaffected though roughly comparable server (i.e. Nathan's 64 core
AMD server). Andres observed that "stalled-cycles-frontend" and
"stalled-cycles-backend" Linux perf events were at huge variance
depending on whether these Intel systems were patched or unpatched.
They were about the same on the AMD system to begin with.

[1] http://www.intel.com/content/dam/www/public/us/en/documents/white-papers/xeon-lock-scaling-analysis-paper.pdf

[Andres Freund]

Hi,

Post by Peter Geoghegan
Here are the results of a benchmark on Nathan Boley's 64-core, 4
socket server: http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/

That's interesting. The maximum number of what you see here (~293125)
is markedly lower than what I can get.

... poke around ...

Hm, that's partially because you're using pgbench without -M prepared if
I see that correctly. The bottleneck in that case is primarily memory
allocation. But even after that I am getting higher
numbers: ~342497.

Trying to nail down the differnce it oddly seems to be your
max_connections=80 vs my 100. The profile in both cases is markedly
different, way much more spinlock contention with 80. All in
Pin/UnpinBuffer().

I think =80 has to lead to some data being badly aligned. I can
reproduce that =91 has *much* better performance than =90. 170841.844938
vs 368490.268577 in a 10s test. Reproducable both with an without the test.
That's certainly worth some investigation.
This is *not* reproducable on the intel machine, so it might the
associativity of the L1/L2 cache on the AMD.

Post by Peter Geoghegan
Perhaps I should have gone past 64 clients, because in the document
"This implies that with oversubscription (more threads running than
available logical CPUs), the performance of spinlocks can depend
heavily on the exact OS scheduler behavior, and may change drastically
with operating system or VM updates."
I haven't bothered with a higher client counts though, because Andres
noted it's the same with 90 clients on this AMD system. Andres: Do you
see big problems when # clients < # logical cores on the affected
Intel systems?

There's some slowdown with the patch applied, but it's not big. Without
it, the slowdown is much earlier.

Post by Peter Geoghegan
There is only a marginal improvement in performance on this big 4
socket system. Andres informs me privately that he has reproduced the
problem on multiple new 4-socket Intel servers, so it seems reasonable
to suppose that more or less an Intel thing.

I've just poked around, it's not just 4 socket, but also 2 socket
systems.

Some background:
The setups that triggered me into working on the patchset didn't really
have a pgbench like workload, the individual queries were/are more
complicated even though it's still an high throughput OLTP workload. And
the contention was *much* higher than what I can reproduce with pgbench
-S, there was often nearly all time spent in the lwlock's spinlock, and
it was primarily the buffer mapping lwlocks, being locked in shared
mode. The difference is that instead of locking very few buffers per
query like pgbench does, they touched much more.
If you look at a profile of a pgbench -S workload -cj64 it's pretty much all
bottlenecked by GetSnapshotData():
unpatched:
- 40.91% postgres_plainl postgres_plainlw [.] s_lock
- s_lock
- 51.34% LWLockAcquire
GetSnapshotData
- GetTransactionSnapshot
+ 55.23% PortalStart
+ 44.77% PostgresMain
- 48.66% LWLockRelease
GetSnapshotData
- GetTransactionSnapshot
+ 53.64% PostgresMain
+ 46.36% PortalStart
+ 2.65% pgbench [kernel.kallsyms] [k] try_to_wake_up
+ 2.61% postgres_plainl postgres_plainlw [.] LWLockRelease
+ 1.36% postgres_plainl postgres_plainlw [.] LWLockAcquire
+ 1.25% postgres_plainl postgres_plainlw [.] GetSnapshotData

patched:
- 2.94% postgres postgres [.] LWLockAcquire
- LWLockAcquire
+ 26.61% ReadBuffer_common
+ 17.08% GetSnapshotData
+ 10.71% _bt_relandgetbuf
+ 10.24% LockAcquireExtended
+ 10.11% VirtualXactLockTableInsert
+ 9.17% VirtualXactLockTableCleanup
+ 7.55% _bt_getbuf
+ 3.40% index_fetch_heap
+ 1.51% LockReleaseAll
+ 0.89% StartTransactionCommand
+ 0.83% _bt_next
+ 0.75% LockRelationOid
+ 0.52% ReadBufferExtended
- 2.75% postgres postgres [.] _bt_compare
- _bt_compare
+ 96.72% _bt_binsrch
+ 1.71% _bt_moveright
+ 1.29% _bt_search
- 2.67% postgres postgres [.] GetSnapshotData
- GetSnapshotData
+ 97.03% GetTransactionSnapshot
+ 2.16% PostgresMain
+ 0.81% PortalStart

So now the profile looks much saner/less contended which immediately is
visible in transaction rates: 192584.218530 vs 552834.002573.

But if you try to attack the contention from the other end, by setting
default_transaction_isolation='repeatable read' to reduce the number of
snapshots taken, its suddenly 536789.807391 vs 566839.328922. A *much*
smaller benefit.
The reason the patch doesn't help that much with that setting is that there
simply isn't as much actual contention there:

+ 2.77% postgres postgres [.] _bt_compare
- 2.72% postgres postgres [.] LWLockAcquire
- LWLockAcquire
+ 30.51% ReadBuffer_common
+ 12.45% VirtualXactLockTableInsert
+ 12.07% _bt_relandgetbuf
+ 10.61% VirtualXactLockTableCleanup
+ 9.95% LockAcquireExtended
+ 8.59% GetSnapshotData
+ 7.51% _bt_getbuf
+ 2.54% index_fetch_heap
+ 1.47% LockReleaseAll
+ 1.16% StartTransactionCommand
+ 0.95% LockRelationOid
+ 0.89% _bt_next
+ 0.77% ReadBufferExtended
- 2.19% postgres postgres [.] s_lock
- s_lock
+ 52.41% PinBuffer
+ 47.04% UnpinBuffer

While locking is still visible, it's not so extraordinarily dominant. If
you then profile using -e stalled-cycles-backend, you can see that the
actual problematic pipeline stalls are caused by Pin/UnpinBuffer. So
that's now the bottleneck. Since that doesn't use lwlocks, it's not
surprising to see an lwlocks patch doesn't bring as much benefit.

Post by Peter Geoghegan
So, I think it's fair to say, given what we now know from Andres'
numbers and the numbers I got from Nathan's server, that this patch is
closer to being something that addresses a particularly unfortunate
pathology on many-socket Intel system than it is to being a general
performance optimization.

I think one would see bigger improvements if we'd ha testcase that also
caused real lock contention in the AMDs. If you look at a profile there:

unpatched, -c default_transaction_isolation=read committed
- 3.64% postgres_plainl postgres_plainlw [.] s_lock
- s_lock
+ 51.49% LWLockAcquire
+ 45.19% LWLockRelease
+ 1.73% PinBuffer
+ 1.59% UnpinBuffer
- 3.46% postgres_plainl postgres_plainlw [.] LWLockAcquire
- LWLockAcquire
+ 36.67% GetSnapshotData
+ 24.61% ReadBuffer_common
+ 11.45% _bt_relandgetbuf
+ 6.78% _bt_getbuf
+ 5.99% LockAcquireExtended
+ 5.27% VirtualXactLockTableInsert
+ 5.00% VirtualXactLockTableCleanup
+ 1.52% _bt_next
+ 1.31% index_fetch_heap
+ 0.76% LockReleaseAll
- 2.90% postgres_plainl postgres_plainlw [.] GetSnapshotData
- GetSnapshotData
+ 97.40% GetTransactionSnapshot
+ 2.31% PostgresMain

unpatched, -c default_transaction_isolation=repeatable read

- 2.78% postgres_plainl postgres_plainlw [.] LWLockAcquire
- LWLockAcquire
+ 35.41% ReadBuffer_common
+ 15.12% _bt_relandgetbuf
+ 11.78% GetSnapshotData
+ 9.29% _bt_getbuf
+ 7.36% VirtualXactLockTableInsert
+ 7.25% LockAcquireExtended
+ 6.67% VirtualXactLockTableCleanup
+ 1.98% index_fetch_heap
+ 1.88% _bt_next
+ 1.47% LockReleaseAll
+ 0.50% StartTransactionCommand
- 2.21% postgres_plainl postgres_plainlw [.] hash_search_with_hash_value
- hash_search_with_hash_value
+ 53.23% BufTableLookup
+ 13.92% LockAcquireExtended
+ 8.99% GetPortalByName
+ 6.79% RelationIdGetRelation
+ 4.98% FetchPreparedStatement
+ 4.43% CreatePortal
+ 3.61% PortalDrop
+ 2.46% RemoveLocalLock

There's obviously not even close as much contention as in the intel
case. Which shows in the benchmark results:
440910.048200 vs. 473789.980486

So, the lwlock patch cannot improve concurrency by the degree it did for
intel, since there simply isn't much of a bottleneck. But I'd be
surprised if there aren't cases with much more prominent bottlenecks,
this just isn't one of them.

The changed algorithm for lwlock imo is an *algorithmic* improvement,
not one for a particular architecture. The advantage being that locking
a lwlock which is primarily taken in shared mode will never need need to
wait or loop.

Post by Peter Geoghegan
Based on the above quoted passage, it isn't
unreasonable to suppose other vendors or architectures could be
affected, but that isn't in evidence. While I welcome the use of
atomic operations in the context of LW_SHARED acquisition as general
progress, I think that to the outside world my proposed messaging is
more useful. It's not quite a bug-fix, but if you're using a
many-socket Intel server, you're *definitely* going to want to use a
PostgreSQL version that is unaffected. You may well not want to take
on the burden of waiting for 9.4, or waiting for it to fully
stabilize.
I note that Andres has a feature branch of this backported to Postgres
9.2, no doubt because of a request from a 2ndQuadrant customer. I have
to wonder if we should think about making this available with a
configure switch in one or more back branches.

Yes, that branch is used by some of them. But to make that clear to all
that are still reading, I have *first* presented the patch & findings to
-hackers and *then* backported it, and I have referenced the existance
of the patch for 9.2 on list before. This isn't some kind of "secret
sauce" deal...

But note that there is one significant difference between the 9.2 and
HEAD version, the former directly uses gcc intrinsics, without having
any form of checks or abstraction across compilers...

Post by Peter Geoghegan
I think that the
complete back-porting of the fsync request queue issue's fix in commit
758728 could be considered a precedent - that too was a fix for a
really bad worst-case that was encountered fairly infrequently in the
wild. It's sort of horrifying to have red-hot spinlocks in production,
so that seems like the kind of thing we should make an effort to
address for those running multi-socket systems. Of those running
Postgres on new multi-socket systems, the reality is that the majority
are running on Intel hardware. Unfortunately, everyone knows that
Intel will soon be the only game in town when it comes to high-end
x86_64 servers, which contributes to my feeling that we need to target
back branches. We should do something about the possible regression
with older compilers using the fallback first, though.

That might be something to do later, as it *really* can hurt in
practice. We had one server go from load 240 to 11...

But I think we should first focus on getting the patch ready for
master, then we can see where it's going. At the very least I'd like to
split of the part modifying the current spinlocks to use the atomics,
that seems far to invasive.

Post by Peter Geoghegan
It would be useful to know more about the nature of the problem that
made such an appreciable difference in Andres' original post.

I unfortunately can't tell you that much more, not because it's private,
but because it mostly was diagnosed by remote hand debugging, limiting
insights considerably.
What I *do* know is that the bottleneck entirely was caused by the
buffer mapping lwlocks, all taken in shared mode according to call graph
profiles. The call graph and the queries I have seen indicated that lots
of the frequent queries involved nested loops over not inconsiderable
number of pages.
We've also tried to increase the number of buffer mapping locks, but
that didn't prove to be very helpful.

Hm, that went on a bit...

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Jeff Janes]

Post by Andres Freund
The setups that triggered me into working on the patchset didn't really
have a pgbench like workload, the individual queries were/are more
complicated even though it's still an high throughput OLTP workload. And
the contention was *much* higher than what I can reproduce with pgbench
-S, there was often nearly all time spent in the lwlock's spinlock, and
it was primarily the buffer mapping lwlocks, being locked in shared
mode. The difference is that instead of locking very few buffers per
query like pgbench does, they touched much more.

Perhaps I should try to argue for this extension to pgbench again:

http://www.postgresql.org/message-id/CAMkU=1w0K3RNhtPuLF8WQoVi6gxgG6mcnpC=-***@mail.gmail.com

I think it would go a good job of exercising what you want, provided you
set the scale so that all data fit in RAM but not in shared_buffers.

Or maybe you want it to fit in shared_buffers, since the buffer mapping
lock was contended in shared mode--that suggests the problem is finding the
buffer that already has the page, not making a buffer to have the page.

Cheers,

Jeff

[Peter Geoghegan]

Post by Andres Freund
The changed algorithm for lwlock imo is an *algorithmic* improvement,
not one for a particular architecture. The advantage being that locking
a lwlock which is primarily taken in shared mode will never need need to
wait or loop.

I agree. My point was only that the messaging ought to be that this is
something that those with multi-socket Intel systems should take note
of.

Post by Andres Freund
Yes, that branch is used by some of them. But to make that clear to all
that are still reading, I have *first* presented the patch & findings to
-hackers and *then* backported it, and I have referenced the existance
of the patch for 9.2 on list before. This isn't some kind of "secret
sauce" deal...

No, of course not. I certainly didn't mean to imply that. My point was
only that anyone that is affected to the same degree as the party with
the 4 socket server might be left with a very poor impression of
Postgres if we failed to fix the problem. It clearly rises to the
level of a bugfix.

Post by Andres Freund
That might be something to do later, as it *really* can hurt in
practice. We had one server go from load 240 to 11...

Well, we have to commit something on master first. But it should be a
priority to avoid having this hurt users further, since the problems
are highly predictable for certain types of servers.

Post by Andres Freund
But I think we should first focus on getting the patch ready for
master, then we can see where it's going. At the very least I'd like to
split of the part modifying the current spinlocks to use the atomics,
that seems far to invasive.

Agreed.

Post by Andres Freund
I unfortunately can't tell you that much more, not because it's private,
but because it mostly was diagnosed by remote hand debugging, limiting
insights considerably.

Of course.

[Peter Geoghegan]

Post by Andres Freund

Post by Peter Geoghegan
Here are the results of a benchmark on Nathan Boley's 64-core, 4
socket server: http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/

That's interesting. The maximum number of what you see here (~293125)
is markedly lower than what I can get.
... poke around ...
Hm, that's partially because you're using pgbench without -M prepared if
I see that correctly. The bottleneck in that case is primarily memory
allocation. But even after that I am getting higher
numbers: ~342497.
Trying to nail down the differnce it oddly seems to be your
max_connections=80 vs my 100. The profile in both cases is markedly
different, way much more spinlock contention with 80. All in
Pin/UnpinBuffer().

I updated this benchmark, with your BufferDescriptors alignment patch
[1] applied on top of master (while still not using "-M prepared" in
order to keep the numbers comparable). So once again, that's:

http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/

It made a bigger, fairly noticeable difference, but not so big a
difference as you describe here. Are you sure that you saw this kind
of difference with only 64 clients, as you mentioned elsewhere [1]
(perhaps you fat-fingered [1] -- "-cj" is ambiguous)? Obviously
max_connections is still 80 in the above. Should I have gone past 64
clients to see the problem? The best numbers I see with the [1] patch
applied on master is only ~327809 for -S 10 64 clients. Perhaps I've
misunderstood.

[1] "Misaligned BufferDescriptors causing major performance problems
on AMD": http://www.postgresql.org/message-id/***@awork2.anarazel.de

[Andres Freund]

Post by Peter Geoghegan

Post by Andres Freund

Post by Peter Geoghegan
Here are the results of a benchmark on Nathan Boley's 64-core, 4
socket server: http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/

That's interesting. The maximum number of what you see here (~293125)
is markedly lower than what I can get.
... poke around ...
Hm, that's partially because you're using pgbench without -M prepared if
I see that correctly. The bottleneck in that case is primarily memory
allocation. But even after that I am getting higher
numbers: ~342497.
Trying to nail down the differnce it oddly seems to be your
max_connections=80 vs my 100. The profile in both cases is markedly
different, way much more spinlock contention with 80. All in
Pin/UnpinBuffer().

I updated this benchmark, with your BufferDescriptors alignment patch
[1] applied on top of master (while still not using "-M prepared" in
http://postgres-benchmarks.s3-website-us-east-1.amazonaws.com/amd-4-socket-rwlocks/
It made a bigger, fairly noticeable difference, but not so big a
difference as you describe here. Are you sure that you saw this kind
of difference with only 64 clients, as you mentioned elsewhere [1]
(perhaps you fat-fingered [1] -- "-cj" is ambiguous)? Obviously
max_connections is still 80 in the above. Should I have gone past 64
clients to see the problem? The best numbers I see with the [1] patch
applied on master is only ~327809 for -S 10 64 clients. Perhaps I've
misunderstood.

That's likely -M prepared. It was with -c 64 -j 64...

Greetings,

Andres Freund

[Heikki Linnakangas]

Post by Andres Freund
Hi,

Post by Peter Geoghegan

Post by Andres Freund
1) I've added an abstracted atomic ops implementation. Needs a fair
amount of work, also submitted as a separate CF entry. (Patch 1 & 2)

Commit 220b34331f77effdb46798ddd7cca0cffc1b2858 caused bitrot when
applying 0002-Very-basic-atomic-ops-implementation.patch. Please
rebase.

I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.
I plan to split the atomics patch into smaller chunks before
reposting. Imo the "Convert the PGPROC->lwWaitLink list into a dlist
instead of open coding it." is worth being applied independently from
the rest of the series, it simplies code and it fixes a bug...

I committed a fix for the WakeupWaiters() bug now, without the rest of
the "open coding" patch. Converting lwWaitLInk into a dlist is probably
a good idea, but seems better to fix the bug separately, for the sake of
git history if nothing else.

- Heikki

[Amit Kapila]

Post by Andres Freund
I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.

As per discussion in developer meeting, I wanted to test shared
buffer scaling patch with this branch. I am getting merge
conflicts as per HEAD. Could you please get it resolved, so that
I can get the data.

Post by Andres Freund
From git://git.postgresql.org/git/users/andresfreund/postgres

* branch rwlock-contention -> FETCH_HEAD
Auto-merging src/test/regress/regress.c
CONFLICT (content): Merge conflict in src/test/regress/regress.c
Auto-merging src/include/storage/proc.h
Auto-merging src/include/storage/lwlock.h
CONFLICT (content): Merge conflict in src/include/storage/lwlock.h
Auto-merging src/include/storage/ipc.h
CONFLICT (content): Merge conflict in src/include/storage/ipc.h
Auto-merging src/include/storage/barrier.h
CONFLICT (content): Merge conflict in src/include/storage/barrier.h
Auto-merging src/include/pg_config_manual.h
Auto-merging src/include/c.h
Auto-merging src/backend/storage/lmgr/spin.c
Auto-merging src/backend/storage/lmgr/proc.c
Auto-merging src/backend/storage/lmgr/lwlock.c
CONFLICT (content): Merge conflict in src/backend/storage/lmgr/lwlock.c
Auto-merging src/backend/storage/ipc/shmem.c
Auto-merging src/backend/storage/ipc/ipci.c
Auto-merging src/backend/access/transam/xlog.c
CONFLICT (content): Merge conflict in src/backend/access/transam/xlog.c
Auto-merging src/backend/access/transam/twophase.c
Auto-merging configure.in
Auto-merging configure
Auto-merging config/c-compiler.m4
Automatic merge failed; fix conflicts and then commit the result.



With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Amit Kapila

Post by Andres Freund
I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.

As per discussion in developer meeting, I wanted to test shared
buffer scaling patch with this branch. I am getting merge
conflicts as per HEAD. Could you please get it resolved, so that
I can get the data.

I have started looking into this patch and have few questions/
findings which are shared below:

1. I think stats for lwstats->ex_acquire_count will be counted twice,
first it is incremented in LWLockAcquireCommon() and then in
LWLockAttemptLock()

2.
Handling of potentialy_spurious case seems to be pending
in LWLock functions like LWLockAcquireCommon().

LWLockAcquireCommon()
{
..
/* add to the queue */
LWLockQueueSelf(l, mode);

/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);

}

I think it can lead to some problems, example:

Session -1
1. Acquire Exclusive LWlock

Session -2
1. Acquire Shared LWlock

1a. Unconditionally incrementing shared count by session-2

Session -1
2. Release Exclusive lock

Session -3
1. Acquire Exclusive LWlock
It will put itself to wait queue by seeing the lock count incremented
by Session-2

Session-2
1b. Decrement the shared count and add it to wait queue.

Session-4
1. Acquire Exclusive lock
This session will get the exclusive lock, because even
though other lockers are waiting, lockcount is zero.

Session-2
2. Try second time to take shared lock, it won't get
as session-4 already has an exclusive lock, so it will
start waiting

Session-4
2. Release Exclusive lock
it will not wake the waiters because waiters have been added
before acquiring this lock.

So in above scenario, Session-3 and Session-2 are waiting in queue
with nobody to awake them.

I have not reproduced the exact scenario above,
so I might be missing some thing which will not
lead to above situation.

3.
LWLockAcquireCommon()
{
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
..
}
proc->lwWaiting is checked, updated without spinklock where
as previously it was done under spinlock, won't it be unsafe?

4.
LWLockAcquireCommon()
{
..
for (;;)
{
/* "false" means cannot accept cancel/die interrupt here. */
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
}

lock->releaseOK is updated/checked without spinklock where
as previously it was done under spinlock, won't it be unsafe?


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Amit Kapila

Post by Andres Freund
I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.

As per discussion in developer meeting, I wanted to test shared
buffer scaling patch with this branch. I am getting merge
conflicts as per HEAD. Could you please get it resolved, so that
I can get the data.

I have started looking into this patch and have few questions/
1. I think stats for lwstats->ex_acquire_count will be counted twice,
first it is incremented in LWLockAcquireCommon() and then in
LWLockAttemptLock()

Hrmpf. Will fix.

Post by Amit Kapila
2.
Handling of potentialy_spurious case seems to be pending
in LWLock functions like LWLockAcquireCommon().
LWLockAcquireCommon()
{
..
/* add to the queue */
LWLockQueueSelf(l, mode);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
}
Session -1
1. Acquire Exclusive LWlock
Session -2
1. Acquire Shared LWlock
1a. Unconditionally incrementing shared count by session-2
Session -1
2. Release Exclusive lock
Session -3
1. Acquire Exclusive LWlock
It will put itself to wait queue by seeing the lock count incremented
by Session-2
Session-2
1b. Decrement the shared count and add it to wait queue.
Session-4
1. Acquire Exclusive lock
This session will get the exclusive lock, because even
though other lockers are waiting, lockcount is zero.
Session-2
2. Try second time to take shared lock, it won't get
as session-4 already has an exclusive lock, so it will
start waiting
Session-4
2. Release Exclusive lock
it will not wake the waiters because waiters have been added
before acquiring this lock.

I don't understand this step here? When releasing the lock it'll notice
that the waiters is <> 0 and acquire the spinlock which should protect
against badness here?

Post by Amit Kapila
3.
LWLockAcquireCommon()
{
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
..
}
proc->lwWaiting is checked, updated without spinklock where
as previously it was done under spinlock, won't it be unsafe?

It was previously checked/unset without a spinlock as well:
/*
* Awaken any waiters I removed from the queue.
*/
while (head != NULL)
{
LOG_LWDEBUG("LWLockRelease", T_NAME(l), T_ID(l), "release waiter");
proc = head;
head = proc->lwWaitLink;
proc->lwWaitLink = NULL;
proc->lwWaiting = false;
PGSemaphoreUnlock(&proc->sem);
}
I don't think there's dangers here, lwWaiting will only ever be
manipulated by the the PGPROC's owner. As discussed elsewhere there
needs to be a write barrier before the proc->lwWaiting = false, even in
upstream code.

Post by Amit Kapila
4.
LWLockAcquireCommon()
{
..
for (;;)
{
/* "false" means cannot accept cancel/die interrupt here. */
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
}
lock->releaseOK is updated/checked without spinklock where
as previously it was done under spinlock, won't it be unsafe?

Hm. That's probably buggy. Good catch. Especially if you have a compiler
that does byte manipulation by reading e.g. 4 bytes from a struct and
then write the wider variable back... So the releaseOk bit needs to move
into LWLockDequeueSelf().

Thanks for looking!

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila

Post by Amit Kapila

Post by Andres Freund
I've pushed a rebased version of the patchset to
http://git.postgresql.org/gitweb/?p=users/andresfreund/postgres.git
branch rwlock contention.
220b34331f77effdb46798ddd7cca0cffc1b2858 actually was the small problem,
ea9df812d8502fff74e7bc37d61bdc7d66d77a7f was the major PITA.

As per discussion in developer meeting, I wanted to test shared
buffer scaling patch with this branch. I am getting merge
conflicts as per HEAD. Could you please get it resolved, so that
I can get the data.

I have started looking into this patch and have few questions/
1. I think stats for lwstats->ex_acquire_count will be counted twice,
first it is incremented in LWLockAcquireCommon() and then in
LWLockAttemptLock()

Hrmpf. Will fix.

Post by Amit Kapila
2.
Handling of potentialy_spurious case seems to be pending
in LWLock functions like LWLockAcquireCommon().
LWLockAcquireCommon()
{
..
/* add to the queue */
LWLockQueueSelf(l, mode);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
}
Session -1
1. Acquire Exclusive LWlock
Session -2
1. Acquire Shared LWlock
1a. Unconditionally incrementing shared count by session-2
Session -1
2. Release Exclusive lock
Session -3
1. Acquire Exclusive LWlock
It will put itself to wait queue by seeing the lock count incremented
by Session-2
Session-2
1b. Decrement the shared count and add it to wait queue.
Session-4
1. Acquire Exclusive lock
This session will get the exclusive lock, because even
though other lockers are waiting, lockcount is zero.
Session-2
2. Try second time to take shared lock, it won't get
as session-4 already has an exclusive lock, so it will
start waiting
Session-4
2. Release Exclusive lock
it will not wake the waiters because waiters have been added
before acquiring this lock.

I don't understand this step here? When releasing the lock it'll notice
that the waiters is <> 0 and acquire the spinlock which should protect
against badness here?

While Releasing lock, I think it will not go to Wakeup waiters
(LWLockWakeup), because releaseOK will be false. releaseOK
can be set as false when Session-1 has Released Exclusive lock
and wakedup some previous waiter. Once it is set to false, it can
be reset to true only for retry logic(after getting semaphore).

Post by Andres Freund

Post by Amit Kapila
3.

I don't think there's dangers here, lwWaiting will only ever be
manipulated by the the PGPROC's owner. As discussed elsewhere there
needs to be a write barrier before the proc->lwWaiting = false, even in
upstream code.

Agreed.


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
2.
Handling of potentialy_spurious case seems to be pending
in LWLock functions like LWLockAcquireCommon().
LWLockAcquireCommon()
{
..
/* add to the queue */
LWLockQueueSelf(l, mode);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
}
Session -1
1. Acquire Exclusive LWlock
Session -2
1. Acquire Shared LWlock
1a. Unconditionally incrementing shared count by session-2
Session -1
2. Release Exclusive lock
Session -3
1. Acquire Exclusive LWlock
It will put itself to wait queue by seeing the lock count incremented
by Session-2
Session-2
1b. Decrement the shared count and add it to wait queue.
Session-4
1. Acquire Exclusive lock
This session will get the exclusive lock, because even
though other lockers are waiting, lockcount is zero.
Session-2
2. Try second time to take shared lock, it won't get
as session-4 already has an exclusive lock, so it will
start waiting
Session-4
2. Release Exclusive lock
it will not wake the waiters because waiters have been added
before acquiring this lock.

I don't understand this step here? When releasing the lock it'll notice
that the waiters is <> 0 and acquire the spinlock which should protect
against badness here?

While Releasing lock, I think it will not go to Wakeup waiters
(LWLockWakeup), because releaseOK will be false. releaseOK
can be set as false when Session-1 has Released Exclusive lock
and wakedup some previous waiter. Once it is set to false, it can
be reset to true only for retry logic(after getting semaphore).

I unfortunately still can't follow. If Session-1 woke up some previous
waiter the woken up process will set releaseOK to true again when it
loops to acquire the lock?


Somewhat unrelated:

I have a fair amount of doubt about the effectiveness of the releaseOK
logic (which imo also is pretty poorly documented).
Essentially its intent is to avoid unneccessary scheduling when other
processes have already been woken up (i.e. releaseOK has been set to
false). I believe the theory is that if any process has already been
woken up it's pointless to wake up additional processes
(i.e. PGSemaphoreUnlock()) because the originally woken up process will
wake up at some point. But if the to-be-woken up process is scheduled
out because it used all his last timeslices fully that means we'll not
wakeup other waiters for a relatively long time.

It's been introduced in the course of
5b9a058384e714b89e050fc0b6381f97037c665a whose logic generally is rather
sound - I just doubt that the releaseOK part is necessary.

It'd certainly interesting to rip releaseOK out and benchmark the
result... My theory is that the average latency will go down on busy
systems that aren't IO bound.

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
2.
Handling of potentialy_spurious case seems to be pending
in LWLock functions like LWLockAcquireCommon().
LWLockAcquireCommon()
{
..
/* add to the queue */
LWLockQueueSelf(l, mode);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
}
Session -1
1. Acquire Exclusive LWlock
Session -2
1. Acquire Shared LWlock
1a. Unconditionally incrementing shared count by session-2
Session -1
2. Release Exclusive lock
Session -3
1. Acquire Exclusive LWlock
It will put itself to wait queue by seeing the lock count incremented
by Session-2
Session-2
1b. Decrement the shared count and add it to wait queue.
Session-4
1. Acquire Exclusive lock
This session will get the exclusive lock, because even
though other lockers are waiting, lockcount is zero.
Session-2
2. Try second time to take shared lock, it won't get
as session-4 already has an exclusive lock, so it will
start waiting
Session-4
2. Release Exclusive lock
it will not wake the waiters because waiters have been added
before acquiring this lock.

I don't understand this step here? When releasing the lock it'll notice
that the waiters is <> 0 and acquire the spinlock which should protect
against badness here?

While Releasing lock, I think it will not go to Wakeup waiters
(LWLockWakeup), because releaseOK will be false. releaseOK
can be set as false when Session-1 has Released Exclusive lock
and wakedup some previous waiter. Once it is set to false, it can
be reset to true only for retry logic(after getting semaphore).

I unfortunately still can't follow.

You have followed it pretty well as far as I can understand from your
replies, as there is no reproducible test (which I think is bit tricky to
prepare), so it becomes difficult to explain by theory.

Post by Andres Freund
If Session-1 woke up some previous
waiter the woken up process will set releaseOK to true again when it
loops to acquire the lock?

You are right, it will wakeup the existing waiters, but I think the
new logic has one difference which is that it can allow the backend to
take Exclusive lock when there are already waiters in queue. As per
above example even though Session-2 and Session-3 are in wait
queue, Session-4 will be able to acquire Exclusive lock which I think
was previously not possible.

Post by Andres Freund
I have a fair amount of doubt about the effectiveness of the releaseOK
logic (which imo also is pretty poorly documented).
Essentially its intent is to avoid unneccessary scheduling when other
processes have already been woken up (i.e. releaseOK has been set to
false). I believe the theory is that if any process has already been
woken up it's pointless to wake up additional processes
(i.e. PGSemaphoreUnlock()) because the originally woken up process will
wake up at some point. But if the to-be-woken up process is scheduled
out because it used all his last timeslices fully that means we'll not
wakeup other waiters for a relatively long time.

I think it will also maintain that the wokedup process won't stall for
very long time, because if we wake new waiters, then previously woked
process can again enter into wait queue and similar thing can repeat
for long time.

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Andres Freund
I unfortunately still can't follow.

You have followed it pretty well as far as I can understand from your
replies, as there is no reproducible test (which I think is bit tricky to
prepare), so it becomes difficult to explain by theory.

I'm working an updated patch that moves the releaseOK into the
spinlocks. Maybe that's the problem already - it's certainly not correct
as is.

Post by Amit Kapila

Post by Andres Freund
If Session-1 woke up some previous
waiter the woken up process will set releaseOK to true again when it
loops to acquire the lock?

You are right, it will wakeup the existing waiters, but I think the
new logic has one difference which is that it can allow the backend to
take Exclusive lock when there are already waiters in queue. As per
above example even though Session-2 and Session-3 are in wait
queue, Session-4 will be able to acquire Exclusive lock which I think
was previously not possible.

I think that was previously possible as well - in a slightly different
set of circumstances though. After a process releases a lock and wakes
up some of several waiters another process can come in and acquire the
lock. Before the woken up process gets scheduled again. lwlocks aren't
fair locks...

Post by Amit Kapila

Post by Andres Freund
I have a fair amount of doubt about the effectiveness of the releaseOK
logic (which imo also is pretty poorly documented).
Essentially its intent is to avoid unneccessary scheduling when other
processes have already been woken up (i.e. releaseOK has been set to
false). I believe the theory is that if any process has already been
woken up it's pointless to wake up additional processes
(i.e. PGSemaphoreUnlock()) because the originally woken up process will
wake up at some point. But if the to-be-woken up process is scheduled
out because it used all his last timeslices fully that means we'll not
wakeup other waiters for a relatively long time.

I think it will also maintain that the wokedup process won't stall for
very long time, because if we wake new waiters, then previously woked
process can again enter into wait queue and similar thing can repeat
for long time.

I don't think it effectively does that - newly incoming lockers ignore
the queue and just acquire the lock. Even if there's some other backend
scheduled to wake up. And shared locks can be acquired when there's
exclusive locks waiting.

I think both are actually critical for performance... Otherwise even a
only lightly contended lock would require scheduler activity when a
processes tries to lock something twice. Given the frequency we acquire
some locks with that'd be disastrous...

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
You have followed it pretty well as far as I can understand from your
replies, as there is no reproducible test (which I think is bit tricky to
prepare), so it becomes difficult to explain by theory.

I'm working an updated patch that moves the releaseOK into the
spinlocks. Maybe that's the problem already - it's certainly not correct
as is.

Sure, I will do the test/performance test with updated patch; you
might want to include some more changes based on comments

Post by Andres Freund

Post by Amit Kapila
You are right, it will wakeup the existing waiters, but I think the
new logic has one difference which is that it can allow the backend to
take Exclusive lock when there are already waiters in queue. As per
above example even though Session-2 and Session-3 are in wait
queue, Session-4 will be able to acquire Exclusive lock which I think
was previously not possible.

I think that was previously possible as well - in a slightly different
set of circumstances though. After a process releases a lock and wakes
up some of several waiters another process can come in and acquire the
lock. Before the woken up process gets scheduled again. lwlocks aren't
fair locks...

Okay, but I think changing behaviour for lwlocks might impact some
tests/applications. As they are not fair, I think defining exact
behaviour is not easy and we don't have any concrete scenario which
can be effected, so there should not be problem in accepting
slightly different behaviour.

Few more comments:

1.
LWLockAcquireCommon()
{
..
iterations++;
}

In current logic, I could not see any use of these *iterations* variable.

2.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
/*
* Somebody else dequeued us and has or will..
..
*/
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
}

Do we want to set result = false; after waking in above code?
The idea behind setting false is to indicate whether we get the lock
immediately or not which previously was decided based on if it needs
to queue itself?

3.
LWLockAcquireCommon()
{
..
/*
* Ok, at this point we couldn't grab the lock on the first try. We
* cannot simply queue ourselves to the end of the list and wait to be
* woken up because by now the lock could long have been released.
* Instead add us to the queue and try to grab the lock again. If we
* suceed we need to revert the queuing and be happy, otherwise we
* recheck the lock. If we still couldn't grab it, we know that the
* other lock will see our queue entries when releasing since they
* existed before we checked for the lock.
*/
/* add to the queue */
LWLockQueueSelf(l, mode);

/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
..
}

a. By reading above code and comments, it is not quite clear why
second attempt is important unless somebody thinks on it or refer
your comments in *Notes* section at top of file. I think it's better to
indicate in some way so that code reader can refer to Notes section or
whereever you are planing to keep those comments.

b. There is typo in above comment suceed/succeed.


4.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
..
}

Setting releaseOK in above context might not be required because if the
control comes in this part of code, it will not retry to acquire another
time.

5.
LWLockWaitForVar()
{
..
else
mustwait = false;

if (!mustwait)
break;
..
}

I think we can directly break in else part in above code.

6.
LWLockWaitForVar()
{
..
/*
* Quick test first to see if it the slot is free right now.
*
* XXX: the caller uses a spinlock before this,...
*/
if (pg_atomic_read_u32(&lock->lockcount) == 0)
return true;
}

Does the part of comment that refers to spinlock is still relevant
after using atomic ops?

7.
LWLockWaitForVar()
{
..
/*
* Add myself to wait queue. Note that this is racy, somebody else
* could wakeup before we're finished queuing.
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);

/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, LW_EXCLUSIVE, false,
&potentially_spurious);
}

Why is it important to Attempt lock after queuing in this case, can't
we just re-check exclusive lock as done before queuing?

8.
LWLockWaitForVar()
{
..
PRINT_LWDEBUG("LWLockAcquire undo queue", lock, mode);
break;
}
else
{
PRINT_LWDEBUG("LWLockAcquire waiting 4", lock, mode);
}
..
}

a. I think instead of LWLockAcquire, here we should use
LWLockWaitForVar
b. Isn't it better to use LOG_LWDEBUG instead of PRINT_LWDEBUG(),
as PRINT_LWDEBUG() is generally used in file at entry of functions to
log info about locks?

9.
LWLockUpdateVar()
{
/* Acquire mutex. Time spent holding mutex should be short! */
SpinLockAcquire(&lock->mutex);
..
}

Current code has an Assert for exclusive lock which is missing in
patch, is there any reason for removing Assert?

10.
LWLockRelease(LWLock *l)
{
..
if (l == held_lwlocks[i].lock)
{
mode = held_lwlocks[i].mode;
..
}

It seems that mode is added to held_lwlocks to use it in LWLockRelease().
If yes, then can we deduce the same from lockcount?

11.
LWLockRelease()
{
..
PRINT_LWDEBUG("LWLockRelease", lock, mode);
}

Shouldn't this be in begining of LWLockRelease function rather than
after processing held_lwlocks array?


12.
#ifdef LWLOCK_DEBUG
lock->owner = MyProc;
#endif

Shouldn't it be reset in LWLockRelease?

13.
* This protects us against both problems from above:
* 1) Nobody can release too quick, before we're queued, since after Phase
2 since we're
* already queued.


Second *since* seems to be typo.

* 2) If somebody spuriously got blocked from acquiring the lock, they will
* get queued in Phase 2 and we can wake them up if neccessary or they
will
* have gotten the lock in Phase 2.

In the above line, I think the second mention of *Phase 2* should be *Phase
3*.

Post by Andres Freund

Post by Amit Kapila
I think it will also maintain that the wokedup process won't stall for
very long time, because if we wake new waiters, then previously woked
process can again enter into wait queue and similar thing can repeat
for long time.

I don't think it effectively does that - newly incoming lockers ignore
the queue and just acquire the lock. Even if there's some other backend
scheduled to wake up. And shared locks can be acquired when there's
exclusive locks waiting.

They ignore the queue but I think they won't wakeup new waiters unless
some previous wokedup waiter again tries to acquire lock (as that will
set releaseOK). I am not sure how much such a restriction helps, but
still I think it reduces the chance of getting it stalled.

Post by Andres Freund
I think both are actually critical for performance... Otherwise even a
only lightly contended lock would require scheduler activity when a
processes tries to lock something twice. Given the frequency we acquire
some locks with that'd be disastrous...

Do you have any suggestion how both behaviours can be retained?

Note - Still there is more to review in this patch, however I feel it is
good idea to start some test/performance test of this patch, if you
agree, then I will start the same with the updated patch (result
of conclusion of current review comments).

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
You have followed it pretty well as far as I can understand from your
replies, as there is no reproducible test (which I think is bit tricky

to

Post by Andres Freund

Post by Amit Kapila
prepare), so it becomes difficult to explain by theory.

I'm working an updated patch that moves the releaseOK into the
spinlocks. Maybe that's the problem already - it's certainly not correct
as is.

Sure, I will do the test/performance test with updated patch; you
might want to include some more changes based on comments

I'm nearly finished in cleaning up the atomics part of the patch which
also includes a bit of cleanup of the lwlocks code.

Post by Amit Kapila
1.
LWLockAcquireCommon()
{
..
iterations++;
}
In current logic, I could not see any use of these *iterations* variable.

It's useful for debugging. Should be gone in the final code.

Post by Amit Kapila
2.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
/*
* Somebody else dequeued us and has or will..
..
*/
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
}
Do we want to set result = false; after waking in above code?
The idea behind setting false is to indicate whether we get the lock
immediately or not which previously was decided based on if it needs
to queue itself?

Hm. I don't think it's clear which version is better.

Post by Amit Kapila
3.
LWLockAcquireCommon()
{
..
/*
* Ok, at this point we couldn't grab the lock on the first try. We
* cannot simply queue ourselves to the end of the list and wait to be
* woken up because by now the lock could long have been released.
* Instead add us to the queue and try to grab the lock again. If we
* suceed we need to revert the queuing and be happy, otherwise we
* recheck the lock. If we still couldn't grab it, we know that the
* other lock will see our queue entries when releasing since they
* existed before we checked for the lock.
*/
/* add to the queue */
LWLockQueueSelf(l, mode);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
..
}
a. By reading above code and comments, it is not quite clear why
second attempt is important unless somebody thinks on it or refer
your comments in *Notes* section at top of file. I think it's better to
indicate in some way so that code reader can refer to Notes section or
whereever you are planing to keep those comments.

Ok.

Post by Amit Kapila
b. There is typo in above comment suceed/succeed.

Thanks, fixed.

Post by Amit Kapila
4.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
..
}
Setting releaseOK in above context might not be required because if the
control comes in this part of code, it will not retry to acquire another
time.

Hm. You're probably right.

Post by Amit Kapila
5.
LWLockWaitForVar()
{
..
else
mustwait = false;
if (!mustwait)
break;
..
}
I think we can directly break in else part in above code.

Well, there's another case of mustwait=false above which is triggered
while the spinlock is held. Don't think it'd get simpler.

Post by Amit Kapila
6.
LWLockWaitForVar()
{
..
/*
* Quick test first to see if it the slot is free right now.
*
* XXX: the caller uses a spinlock before this,...
*/
if (pg_atomic_read_u32(&lock->lockcount) == 0)
return true;
}
Does the part of comment that refers to spinlock is still relevant
after using atomic ops?

Yes. pg_atomic_read_u32() isn't a memory barrier (and explicitly
documented not to be).

Post by Amit Kapila
7.
LWLockWaitForVar()
{
..
/*
* Add myself to wait queue. Note that this is racy, somebody else
* could wakeup before we're finished queuing.
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, LW_EXCLUSIVE, false,
&potentially_spurious);
}
Why is it important to Attempt lock after queuing in this case, can't
we just re-check exclusive lock as done before queuing?

Well, that's how Heikki designed LWLockWaitForVar().

Post by Amit Kapila
8.
LWLockWaitForVar()
{
..
PRINT_LWDEBUG("LWLockAcquire undo queue", lock, mode);
break;
}
else
{
PRINT_LWDEBUG("LWLockAcquire waiting 4", lock, mode);
}
..
}
a. I think instead of LWLockAcquire, here we should use
LWLockWaitForVar

right.

Post by Amit Kapila
b. Isn't it better to use LOG_LWDEBUG instead of PRINT_LWDEBUG(),
as PRINT_LWDEBUG() is generally used in file at entry of functions to
log info about locks?

Fine with me.

Post by Amit Kapila
9.
LWLockUpdateVar()
{
/* Acquire mutex. Time spent holding mutex should be short! */
SpinLockAcquire(&lock->mutex);
..
}
Current code has an Assert for exclusive lock which is missing in
patch, is there any reason for removing Assert?

That assert didn't use to be there on master... I'll add it again.

Post by Amit Kapila
10.
LWLockRelease(LWLock *l)
{
..
if (l == held_lwlocks[i].lock)
{
mode = held_lwlocks[i].mode;
..
}
It seems that mode is added to held_lwlocks to use it in LWLockRelease().
If yes, then can we deduce the same from lockcount?

No. It can be temporarily too high (the whole backout stuff). It's also
much cheaper to test a process local variable.

Post by Amit Kapila
11.
LWLockRelease()
{
..
PRINT_LWDEBUG("LWLockRelease", lock, mode);
}
Shouldn't this be in begining of LWLockRelease function rather than
after processing held_lwlocks array?

Ok.

Post by Amit Kapila
12.
#ifdef LWLOCK_DEBUG
lock->owner = MyProc;
#endif
Shouldn't it be reset in LWLockRelease?

That's actually intentional. It's quite useful to know the last owner
when debugging lwlock code.

Post by Amit Kapila
* 2) If somebody spuriously got blocked from acquiring the lock, they will
* get queued in Phase 2 and we can wake them up if neccessary or they
will
* have gotten the lock in Phase 2.
In the above line, I think the second mention of *Phase 2* should be *Phase
3*.

Right, good catch.

Post by Amit Kapila

Post by Andres Freund
I think both are actually critical for performance... Otherwise even a
only lightly contended lock would require scheduler activity when a
processes tries to lock something twice. Given the frequency we acquire
some locks with that'd be disastrous...

Do you have any suggestion how both behaviours can be retained?

Not sure what you mean. They currently *are* retained? Or do you mean
whether they could be retained while making lwlocks fai?

Post by Amit Kapila
Note - Still there is more to review in this patch, however I feel it is
good idea to start some test/performance test of this patch, if you
agree, then I will start the same with the updated patch (result
of conclusion of current review comments).

I'll post a new version of this + the atomics patch tomorrow. Since the
whole atomics stuff has changed noticeably it probably makes sense to
wait till then.

Thanks for the look!

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
2.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
/*
* Somebody else dequeued us and has or will..
..
*/
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
}
Do we want to set result = false; after waking in above code?
The idea behind setting false is to indicate whether we get the lock
immediately or not which previously was decided based on if it needs
to queue itself?

Hm. I don't think it's clear which version is better.

I thought if we get the lock at first attempt, then result should be
true which seems to be clear, but for the case of second attempt you
are right that it's not clear. In such a case, I think we can go either
way and then later during tests or otherwise if any problem is discovered,
we can revert it.

Post by Andres Freund

Post by Amit Kapila
7.
LWLockWaitForVar()
{
..
/*
* Add myself to wait queue. Note that this is racy, somebody else
* could wakeup before we're finished queuing.
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, LW_EXCLUSIVE, false,
&potentially_spurious);
}
Why is it important to Attempt lock after queuing in this case, can't
we just re-check exclusive lock as done before queuing?

Well, that's how Heikki designed LWLockWaitForVar().

In that case I might be missing some point here, un-patched code of
LWLockWaitForVar() never tries to acquire the lock, but the new code
does so. Basically I am not able to think what is the problem if we just
do below after queuing:
mustwait = pg_atomic_read_u32(&lock->lockcount) != 0;

Could you please explain what is the problem in just rechecking?

Post by Andres Freund

Post by Amit Kapila

Post by Andres Freund
I think both are actually critical for performance... Otherwise even a
only lightly contended lock would require scheduler activity when a
processes tries to lock something twice. Given the frequency we acquire
some locks with that'd be disastrous...

Do you have any suggestion how both behaviours can be retained?

Not sure what you mean.

I just wanted to say that current behaviour of releaseOK seems to
be of use for some cases and if you want to change it, then would it
retain the current behaviour we get by releaseOK?

I understand that till now your patch has not changed anything specific
to releaseOK, but by above discussion I got the impression that you are
planing to change it, that's why I had asked above question.

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
7.
LWLockWaitForVar()
{
..
/*
* Add myself to wait queue. Note that this is racy, somebody else
* could wakeup before we're finished queuing.
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);
/* we're now guaranteed to be woken up if necessary */
mustwait = LWLockAttemptLock(l, LW_EXCLUSIVE, false,
&potentially_spurious);
}
Why is it important to Attempt lock after queuing in this case, can't
we just re-check exclusive lock as done before queuing?

Well, that's how Heikki designed LWLockWaitForVar().

In that case I might be missing some point here, un-patched code of
LWLockWaitForVar() never tries to acquire the lock, but the new code
does so. Basically I am not able to think what is the problem if we just
mustwait = pg_atomic_read_u32(&lock->lockcount) != 0;
Could you please explain what is the problem in just rechecking?

I have further reviewed the lwlock related changes and thought
its good to share my findings with you. This completes my initial
review for lwlock related changes and below are my findings:

1.
LWLockRelease()
{
..
TRACE_POSTGRESQL_LWLOCK_RELEASE(T_NAME(l), T_ID(l));
}

Dynamic tracing macro seems to be omitted from LWLockRelease()
call.

2.
LWLockWakeup()
{
..
#ifdef LWLOCK_STATS
lwstats->spin_delay_count += SpinLockAcquire(&lock->mutex);
#else
SpinLockAcquire(&lock->mutex);
#endif
..
}

Earlier while releasing lock, we don't count it towards LWLock stats
spin_delay_count. I think if we see other places in lwlock.c, it only gets
counted when we try to acquire it in a loop.

3.
LWLockRelease()
{
..
/* grant permission to run, even if a spurious share lock increases
lockcount */
else if (mode == LW_EXCLUSIVE && have_waiters)
check_waiters = true;
/* nobody has this locked anymore, potential exclusive lockers get a chance
*/
else if (lockcount == 0 && have_waiters)
check_waiters = true;
..
}

It seems comments have been reversed in above code.

4.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
..
}

Shouldn't we need to use volatile variable in above loop (lock instead of
l)?

5.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &wakeup)
{
PGPROC *waiter = dlist_container(PGPROC, lwWaitLink, iter.cur);
LOG_LWDEBUG("LWLockRelease", l, mode, "release waiter");
dlist_delete(&waiter->lwWaitLink);
pg_write_barrier();
waiter->lwWaiting = false;
PGSemaphoreUnlock(&waiter->sem);
}
..
}

Why can't we decrement the nwaiters after waking up? I don't think
there is any major problem even if callers do that themselves, but
in some rare cases LWLockRelease() might spuriously assume that
there are some waiters and tries to call LWLockWakeup(). Although
this doesn't create any problem, keeping the value sane is good unless
there is some problem in doing so.

6.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
if (wokeup_somebody && waiter->lwWaitMode == LW_EXCLUSIVE)
continue;
..
if (waiter->lwWaitMode != LW_WAIT_UNTIL_FREE)
{
..
wokeup_somebody = true;
}
..
}
..
}

a.
IIUC above logic, if the waiter queue is as follows:
(S-Shared; X-Exclusive) S X S S S X S S

it can skip the exclusive waiters and release shared waiter.

If my understanding is right, then I think instead of continue, there
should be *break* in above logic.

b.
Consider below sequence of waiters:
(S-Shared; X-Exclusive) S S X S S

I think as per un-patched code, it will wakeup waiters uptill (including)
first Exclusive, but patch will wake up uptill (*excluding*) first
Exclusive.

7.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
dlist_delete(&waiter->lwWaitLink);
dlist_push_tail(&wakeup, &waiter->lwWaitLink);
..
}
..
}

Use of dlist has simplified the code, but I think there might be a slight
overhead of maintaining wakeup queue as compare to un-patched
mechanism especially when there is a long waiter queue.

8.
LWLockConditionalAcquire()
{
..
/*
* We ran into an exclusive lock and might have blocked another
* exclusive lock from taking a shot because it took a time to back
* off. Retry till we are either sure we didn't block somebody (because
* somebody else certainly has the lock) or till we got it.
*
* We cannot rely on the two-step lock-acquisition protocol as in
* LWLockAcquire because we're not using it.
*/
if (potentially_spurious)
{
SPIN_DELAY();
goto retry;
}
..
}

Due to above logic, I think it can keep on retrying for long time before
it actually concludes whether it got lock or not incase other backend/'s
takes Exclusive lock after *double_check* and release before
unconditional increment of shared lock in function LWLockAttemptLock.
I understand that it might be difficult to have such a practical scenario,
however still there is a theoratical possibility of same.

Is there any advantage of retrying in LWLockConditionalAcquire()?

I think its improtant to have 2-phase LockAttempt in case of
LWLockAcquireCommon() as we have splitted the work of trying to
acquire a lock and queuing it for wait incase didn't got the lock,
but here there is no such thing, so I am wondering is there any
problem, if we just return false after failure of first attempt?

9.
LWLockAcquireOrWait()
{
..
/*
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious_first);

if (mustwait)
{
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);

mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious_second);

}

In this function, it doesn't seem to be required to use the return value
*mustwait* of second LWLockAttemptLock() call as a return value of
function, as per usage of this function if we don't get the lock at first
attempt, then it needs to check if the corresponding WAL record is flushed.
I think here we need the logic what you have used in LWLockWaitForVar()
(release the lock if we get it in second attempt).

10.
LWLockAcquireOrWait()
{
..
bool potentially_spurious_first;
bool potentially_spurious_second;
..
}

Why to use *_first and *_second in this function, can't we just have
one variable as in other LWLock.. variant functions?

11.
LWLockAcquireOrWait()
{
..
Assert(mode == LW_SHARED || mode == LW_EXCLUSIVE);
}

Isn't it better to use AssertArg() rather than Assert in above usage?

12.
LWLockAcquireOrWait()
{
..
if (mustwait)
{
/*
* Wait until awakened. Like in LWLockAcquire, be prepared for bogus
* wakups, because we share the semaphore with ProcWaitForSignal.
*/
LOG_LWDEBUG("LWLockAcquireOrWait", T_NAME(l), T_ID(l), "waiting");
..
}

log for awakened is missing, it's there is current code.

13.
LWLockAcquireOrWait()
{
..
LOG_LWDEBUG("LWLockAcquireOrWait", lockid, mode, "suceeded");
..
}

a. spelling of "suceeded" is wrong.
b. such a log is not there in any other LWLock.. variants, if we want
to introduce it, then shouldn't it be done at other places as well.

14.
typedef struct lwlock_stats
{
+ int ex_race;
..
}

Currently I don't see the usage of this variable, is there a plan to use
it in future?

15.
/* must be greater than MAX_BACKENDS */
#define SHARED_LOCK_MASK (~EXCLUSIVE_LOCK)

a. how can we guarantee it to be greater than MaxBackends,
as MaxBackends is int (the max value for which will be
equal to SHARED_LOCK_MASK)?
b. This is used only for LWLOCK_STATS, so shouldn't we
define this under LWLOCK_STATS.

16.
LWLockAcquireCommon()
{
volatile LWLock *lock = l;
..
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious);
..
lock->releaseOK = true;
..
pg_atomic_fetch_sub_u32(&lock->nwaiters, 1);
}

Shouldn't we need to use *lock* variable in LWLockAttemptLock()
function call?


Note - I have completed the review of LWLock related changes of your
overall patch in 3 different parts, as the changes are more and it makes
me understand your views behind implementation. I am maintaining
all the findings, so when you send the updated patch, I will verify by
using the same. I hope that it's not inconvenient for you.

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila
2.
LWLockWakeup()
{
..
#ifdef LWLOCK_STATS
lwstats->spin_delay_count += SpinLockAcquire(&lock->mutex);
#else
SpinLockAcquire(&lock->mutex);
#endif
..
}
Earlier while releasing lock, we don't count it towards LWLock stats
spin_delay_count. I think if we see other places in lwlock.c, it only gets
counted when we try to acquire it in a loop.

I think the previous situation was clearly suboptimal. I've now modified
things so all spinlock acquirations are counted.

Post by Amit Kapila
3.
LWLockRelease()
{
..
/* grant permission to run, even if a spurious share lock increases
lockcount */
else if (mode == LW_EXCLUSIVE && have_waiters)
check_waiters = true;
/* nobody has this locked anymore, potential exclusive lockers get a chance
*/
else if (lockcount == 0 && have_waiters)
check_waiters = true;
..
}
It seems comments have been reversed in above code.

No, they look right. But I've expanded them in the version I'm going to
post in a couple minutes.

Post by Amit Kapila
5.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &wakeup)
{
PGPROC *waiter = dlist_container(PGPROC, lwWaitLink, iter.cur);
LOG_LWDEBUG("LWLockRelease", l, mode, "release waiter");
dlist_delete(&waiter->lwWaitLink);
pg_write_barrier();
waiter->lwWaiting = false;
PGSemaphoreUnlock(&waiter->sem);
}
..
}
Why can't we decrement the nwaiters after waking up? I don't think
there is any major problem even if callers do that themselves, but
in some rare cases LWLockRelease() might spuriously assume that
there are some waiters and tries to call LWLockWakeup(). Although
this doesn't create any problem, keeping the value sane is good unless
there is some problem in doing so.

That won't work because then LWLockWakeup() wouldn't be called when
necessary - precisely because nwaiters is 0.

Post by Amit Kapila
6.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
if (wokeup_somebody && waiter->lwWaitMode == LW_EXCLUSIVE)
continue;
..
if (waiter->lwWaitMode != LW_WAIT_UNTIL_FREE)
{
..
wokeup_somebody = true;
}
..
}
..
}
a.
(S-Shared; X-Exclusive) S X S S S X S S
it can skip the exclusive waiters and release shared waiter.
If my understanding is right, then I think instead of continue, there
should be *break* in above logic.

No, it looks correct to me. What happened is that the first S was woken
up. So there's no point in waking up an exclusive locker, but further
non-exclusive lockers can be woken up.

Post by Amit Kapila
b.
(S-Shared; X-Exclusive) S S X S S
I think as per un-patched code, it will wakeup waiters uptill (including)
first Exclusive, but patch will wake up uptill (*excluding*) first
Exclusive.

I don't think the current code does that. And it'd be a pretty pointless
behaviour, leading to useless increased contention. The only time it'd
make sense for X to be woken up is when it gets run faster than the S
processes.

Post by Amit Kapila
7.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
dlist_delete(&waiter->lwWaitLink);
dlist_push_tail(&wakeup, &waiter->lwWaitLink);
..
}
..
}
Use of dlist has simplified the code, but I think there might be a slight
overhead of maintaining wakeup queue as compare to un-patched
mechanism especially when there is a long waiter queue.

I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.
And the code is *so* much more readable.

Post by Amit Kapila
8.
LWLockConditionalAcquire()
{
..
/*
* We ran into an exclusive lock and might have blocked another
* exclusive lock from taking a shot because it took a time to back
* off. Retry till we are either sure we didn't block somebody (because
* somebody else certainly has the lock) or till we got it.
*
* We cannot rely on the two-step lock-acquisition protocol as in
* LWLockAcquire because we're not using it.
*/
if (potentially_spurious)
{
SPIN_DELAY();
goto retry;
}
..
}
Due to above logic, I think it can keep on retrying for long time before
it actually concludes whether it got lock or not incase other backend/'s
takes Exclusive lock after *double_check* and release before
unconditional increment of shared lock in function LWLockAttemptLock.
I understand that it might be difficult to have such a practical scenario,
however still there is a theoratical possibility of same.

I'm not particularly concerned. We could optimize it a bit, but I really
don't think it's necessary.

Post by Amit Kapila
Is there any advantage of retrying in LWLockConditionalAcquire()?

It's required for correctness. We only retry if we potentially blocked
an exclusive acquirer (by spuriously incrementing/decrementing lockcount
with 1). We need to be sure to either get the lock (in which case we can
wake up the waiter on release), or be sure that we didn't disturb
anyone.

Post by Amit Kapila
9.
LWLockAcquireOrWait()
{
..
/*
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious_first);
if (mustwait)
{
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);
mustwait = LWLockAttemptLock(l, mode, false, &potentially_spurious_second);
}
In this function, it doesn't seem to be required to use the return value
*mustwait* of second LWLockAttemptLock() call as a return value of
function, as per usage of this function if we don't get the lock at first
attempt, then it needs to check if the corresponding WAL record is flushed.

I don't think that's the appropriate comparison. Acquiring the lock in
the lock in the new implementation essentially consists out of these two
steps. We *DID* get the lock here. Without sleeping. So returning the
appropriate return code is correct.
In fact, returning false would break things, because the caller would
hold the lock without freeing it again?

Post by Amit Kapila
11.
LWLockAcquireOrWait()
{
..
Assert(mode == LW_SHARED || mode == LW_EXCLUSIVE);
}
Isn't it better to use AssertArg() rather than Assert in above usage?

I've changed that. But I have to admit, I don't really see the point of
AssertArg(). If it'd output the value, then it'd be beneficial, but it
doesn't.

Post by Amit Kapila
15.
/* must be greater than MAX_BACKENDS */
#define SHARED_LOCK_MASK (~EXCLUSIVE_LOCK)
a. how can we guarantee it to be greater than MaxBackends,
as MaxBackends is int (the max value for which will be
equal to SHARED_LOCK_MASK)?

MaxBackends luckily is limited to something lower. I've added a comment
to that regard.

Post by Amit Kapila
b. This is used only for LWLOCK_STATS, so shouldn't we
define this under LWLOCK_STATS.

It's a general value, so I don't think that's appropriate.

Thanks for the review!

Greetings,

Andres Freund

[Andres Freund]

Post by Andres Freund

Post by Amit Kapila
5.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &wakeup)
{
PGPROC *waiter = dlist_container(PGPROC, lwWaitLink, iter.cur);
LOG_LWDEBUG("LWLockRelease", l, mode, "release waiter");
dlist_delete(&waiter->lwWaitLink);
pg_write_barrier();
waiter->lwWaiting = false;
PGSemaphoreUnlock(&waiter->sem);
}
..
}
Why can't we decrement the nwaiters after waking up? I don't think
there is any major problem even if callers do that themselves, but
in some rare cases LWLockRelease() might spuriously assume that
there are some waiters and tries to call LWLockWakeup(). Although
this doesn't create any problem, keeping the value sane is good unless
there is some problem in doing so.

That won't work because then LWLockWakeup() wouldn't be called when
necessary - precisely because nwaiters is 0.

Err, this is bogus. Memory fail.

The reason I've done so is that it's otherwise much harder to debug
issues where there are backend that have been woken up already, but
haven't rerun yet. Without this there's simply no evidence of that
state. I can't see this being relevant for performance, so I'd rather
have it stay that way.

Greetings,

Andres Freund

[Robert Haas]

Post by Andres Freund
I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.
And the code is *so* much more readable.

I find the slist/dlist stuff actually quite difficult to get right
compared to a hand-rolled linked list. But the really big problem is
that the debugger can't do anything useful with it. You have to work
out the structure-member offset in order to walk the list and manually
cast to char *, adjust the pointer, and cast back. That sucks.
--
Robert Haas
EnterpriseDB: http://www.enterprisedb.com
The Enterprise PostgreSQL Company

[Andres Freund]

Post by Robert Haas

Post by Andres Freund
I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.
And the code is *so* much more readable.

I find the slist/dlist stuff actually quite difficult to get right
compared to a hand-rolled linked list.

Really? I've spent more than a day debugging things with the current
code. And Heikki introduced a bug in it. If you look at how the code
looks before/after I find the difference pretty clear.

Post by Robert Haas
But the really big problem is
that the debugger can't do anything useful with it. You have to work
out the structure-member offset in order to walk the list and manually
cast to char *, adjust the pointer, and cast back. That sucks.

Hm. I can just do that with the debugger here. Not sure if that's
because I added the right thing to my .gdbinit or because I use the
correct compiler flags.

Greetings,

Andres Freund

[Alvaro Herrera]

Post by Robert Haas

Post by Andres Freund
I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.
And the code is *so* much more readable.

I find the slist/dlist stuff actually quite difficult to get right
compared to a hand-rolled linked list. But the really big problem is
that the debugger can't do anything useful with it. You have to work
out the structure-member offset in order to walk the list and manually
cast to char *, adjust the pointer, and cast back. That sucks.

As far as I recall you can get gdb to understand those pointer games
by defining some structs or macros. Maybe we can improve by documenting
this.

[Andres Freund]

Post by Alvaro Herrera

Post by Robert Haas

Post by Andres Freund
I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.
And the code is *so* much more readable.

I find the slist/dlist stuff actually quite difficult to get right
compared to a hand-rolled linked list. But the really big problem is
that the debugger can't do anything useful with it. You have to work
out the structure-member offset in order to walk the list and manually
cast to char *, adjust the pointer, and cast back. That sucks.

As far as I recall you can get gdb to understand those pointer games
by defining some structs or macros. Maybe we can improve by documenting
this.

So, what makes it work for me (among other unrelated stuff) seems to be
the following in .gdbinit, defineing away some things that gdb doesn't
handle:
macro define __builtin_offsetof(T, F) ((int) &(((T *) 0)->F))
macro define __extension__
macro define AssertVariableIsOfTypeMacro(x, y) ((void)0)

Additionally I have "-ggdb -g3" in CFLAGS. That way gdb knows about
postgres' macros. At least if you're in the right scope.

As an example, the following works:
(gdb) p dlist_is_empty(&BackendList) ? NULL : dlist_head_element(Backend, elem, &BackendList)

Greetings,

Andres Freund

[Robert Haas]

Post by Andres Freund
So, what makes it work for me (among other unrelated stuff) seems to be
the following in .gdbinit, defineing away some things that gdb doesn't
macro define __builtin_offsetof(T, F) ((int) &(((T *) 0)->F))
macro define __extension__
macro define AssertVariableIsOfTypeMacro(x, y) ((void)0)
Additionally I have "-ggdb -g3" in CFLAGS. That way gdb knows about
postgres' macros. At least if you're in the right scope.
(gdb) p dlist_is_empty(&BackendList) ? NULL : dlist_head_element(Backend, elem, &BackendList)

Ah, cool. I'll try that.
--
Robert Haas
EnterpriseDB: http://www.enterprisedb.com
The Enterprise PostgreSQL Company

[Andres Freund]

Post by Robert Haas

Post by Andres Freund
So, what makes it work for me (among other unrelated stuff) seems to be
the following in .gdbinit, defineing away some things that gdb doesn't
macro define __builtin_offsetof(T, F) ((int) &(((T *) 0)->F))
macro define __extension__
macro define AssertVariableIsOfTypeMacro(x, y) ((void)0)
Additionally I have "-ggdb -g3" in CFLAGS. That way gdb knows about
postgres' macros. At least if you're in the right scope.
(gdb) p dlist_is_empty(&BackendList) ? NULL : dlist_head_element(Backend, elem, &BackendList)

Ah, cool. I'll try that.

If that works for you, should we put it somewhere in the docs? If so,
where?

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
2.
LWLockWakeup()
{
..
#ifdef LWLOCK_STATS
lwstats->spin_delay_count += SpinLockAcquire(&lock->mutex);
#else
SpinLockAcquire(&lock->mutex);
#endif
..
}
Earlier while releasing lock, we don't count it towards LWLock stats
spin_delay_count. I think if we see other places in lwlock.c, it only gets
counted when we try to acquire it in a loop.

I think the previous situation was clearly suboptimal. I've now modified
things so all spinlock acquirations are counted.

Code has mainly 4 stats (sh_acquire_count, ex_acquire_count,
block_count, spin_delay_count) to track, if I try to see
all stats together to understand the contention situation,
the unpatched code makes sense. spin_delay_count gives
how much delay has happened to acquire spinlock which when
combined with other stats gives the clear situation about
the contention around aquisation of corresponding LWLock.
Now if we want to count the spin lock delay for Release call
as well, then the meaning of the stat is getting changed.
It might be that new meaning of spin_delay_count stat is more
useful in some situations, however the older one has its own
benefits, so I am not sure if changing this as part of this
patch is the best thing to do.

Post by Andres Freund

Post by Amit Kapila
3.
LWLockRelease()
{
..
/* grant permission to run, even if a spurious share lock increases
lockcount */
else if (mode == LW_EXCLUSIVE && have_waiters)
check_waiters = true;
/* nobody has this locked anymore, potential exclusive lockers get a chance
*/
else if (lockcount == 0 && have_waiters)
check_waiters = true;
..
}
It seems comments have been reversed in above code.

No, they look right. But I've expanded them in the version I'm going to
post in a couple minutes.

okay.

Post by Andres Freund

Post by Amit Kapila
5.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &wakeup)
{
PGPROC *waiter = dlist_container(PGPROC, lwWaitLink, iter.cur);
LOG_LWDEBUG("LWLockRelease", l, mode, "release waiter");
dlist_delete(&waiter->lwWaitLink);
pg_write_barrier();
waiter->lwWaiting = false;
PGSemaphoreUnlock(&waiter->sem);
}
..
}
Why can't we decrement the nwaiters after waking up? I don't think
there is any major problem even if callers do that themselves, but
in some rare cases LWLockRelease() might spuriously assume that
there are some waiters and tries to call LWLockWakeup(). Although
this doesn't create any problem, keeping the value sane is good unless
there is some problem in doing so.

That won't work because then LWLockWakeup() wouldn't be called when
necessary - precisely because nwaiters is 0.
The reason I've done so is that it's otherwise much harder to debug
issues where there are backend that have been woken up already, but
haven't rerun yet. Without this there's simply no evidence of that
state. I can't see this being relevant for performance, so I'd rather
have it stay that way.

I am not sure what useful information we can get during debugging by not
doing this in LWLockWakeup() and w.r.t performance it can lead extra
function call, few checks and I think in some cases even can
acquire/release spinlock.

Post by Andres Freund

Post by Amit Kapila
6.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
if (wokeup_somebody && waiter->lwWaitMode == LW_EXCLUSIVE)
continue;
..
if (waiter->lwWaitMode != LW_WAIT_UNTIL_FREE)
{
..
wokeup_somebody = true;
}
..
}
..
}
a.
(S-Shared; X-Exclusive) S X S S S X S S
it can skip the exclusive waiters and release shared waiter.
If my understanding is right, then I think instead of continue, there
should be *break* in above logic.

No, it looks correct to me. What happened is that the first S was woken
up. So there's no point in waking up an exclusive locker, but further
non-exclusive lockers can be woken up.

Okay, even then it makes the current logic of wakingup
different which I am not sure is what this patch is intended
for.

Post by Andres Freund

Post by Amit Kapila
b.
(S-Shared; X-Exclusive) S S X S S
I think as per un-patched code, it will wakeup waiters uptill (including)
first Exclusive, but patch will wake up uptill (*excluding*) first
Exclusive.

I don't think the current code does that.

LWLockRelease()
{
..
/*
* If the front waiter wants exclusive lock, awaken him only.
*
Otherwise awaken as many waiters as want shared access.
*/
if (proc-

Post by Andres Freund
lwWaitMode != LW_EXCLUSIVE)

{
while (proc->lwWaitLink !=
NULL &&
proc->lwWaitLink->lwWaitMode != LW_EXCLUSIVE)
{
if (proc->lwWaitMode != LW_WAIT_UNTIL_FREE)
releaseOK = false;
proc = proc->lwWaitLink;
}
}
/* proc is now the last PGPROC to be
released */
lock->head = proc->lwWaitLink;
proc->lwWaitLink = NULL;
..
}

In the above code, if the first waiter to be woken up is Exclusive waiter,
then it will woke that waiter, else shared waiters till it got
the first exclusive waiter and then first exlusive waiter.

Post by Andres Freund
And it'd be a pretty pointless
behaviour, leading to useless increased contention. The only time it'd
make sense for X to be woken up is when it gets run faster than the S
processes.

Do we get any major benefit by changing the logic of waking up waiters?

Post by Andres Freund

Post by Amit Kapila
7.
LWLockWakeup()
{
..
dlist_foreach_modify(iter, (dlist_head *) &l->waiters)
{
..
dlist_delete(&waiter->lwWaitLink);
dlist_push_tail(&wakeup, &waiter->lwWaitLink);
..
}
..
}
Use of dlist has simplified the code, but I think there might be a slight
overhead of maintaining wakeup queue as compare to un-patched
mechanism especially when there is a long waiter queue.

I don't see that as being relevant. The difference is an instruction or
two - in the slow path we'll enter the kernel and sleep. This doesn't
matter in comparison.

Okay, however I see Robert has also raised a point on this issue
which I am not sure is concluded.

Post by Andres Freund
And the code is *so* much more readable.

Code is more readable, but I don't understand why you
want to do refactoring as part of this patch which ideally
doesn't get any benefit from the same.

Post by Andres Freund

Post by Amit Kapila
8.
LWLockConditionalAcquire()
{
..
/*
* We ran into an exclusive lock and might have blocked another
* exclusive lock from taking a shot because it took a time to back
* off. Retry till we are either sure we didn't block somebody (because
* somebody else certainly has the lock) or till we got it.
*
* We cannot rely on the two-step lock-acquisition protocol as in
* LWLockAcquire because we're not using it.
*/
if (potentially_spurious)
{
SPIN_DELAY();
goto retry;
}
..
}
Due to above logic, I think it can keep on retrying for long time before
it actually concludes whether it got lock or not incase other backend/'s
takes Exclusive lock after *double_check* and release before
unconditional increment of shared lock in function LWLockAttemptLock.
I understand that it might be difficult to have such a practical scenario,
however still there is a theoratical possibility of same.

I'm not particularly concerned. We could optimize it a bit, but I really
don't think it's necessary.

No issues, I was slightly worried about cases where this
API wasn't suppose to take time earlier (like for contentlock in
BufferAlloc), but now it starts taking time. I am not able to see
any specific issue, so let's proceed with keeping the code
as you have in patch and will optimize later incase we face any problem.

Post by Andres Freund

Post by Amit Kapila
Is there any advantage of retrying in LWLockConditionalAcquire()?

It's required for correctness. We only retry if we potentially blocked
an exclusive acquirer (by spuriously incrementing/decrementing lockcount
with 1). We need to be sure to either get the lock (in which case we can
wake up the waiter on release), or be sure that we didn't disturb
anyone.

Okay, got the point.

Post by Andres Freund

Post by Amit Kapila
9.
LWLockAcquireOrWait()
{
..
/*
* NB: We're using nearly the same twice-in-a-row lock acquisition
* protocol as LWLockAcquire(). Check its comments for details.
*/
mustwait = LWLockAttemptLock(l, mode, false,

&potentially_spurious_first);

Post by Andres Freund

Post by Amit Kapila
if (mustwait)
{
LWLockQueueSelf(l, LW_WAIT_UNTIL_FREE);
mustwait = LWLockAttemptLock(l, mode, false,

&potentially_spurious_second);

Post by Andres Freund

Post by Amit Kapila
}
In this function, it doesn't seem to be required to use the return value
*mustwait* of second LWLockAttemptLock() call as a return value of
function, as per usage of this function if we don't get the lock at first
attempt, then it needs to check if the corresponding WAL record is flushed.

I don't think that's the appropriate comparison. Acquiring the lock in
the lock in the new implementation essentially consists out of these two
steps. We *DID* get the lock here. Without sleeping. So returning the
appropriate return code is correct.
In fact, returning false would break things, because the caller would
hold the lock without freeing it again?

Okay, I understand this point, but OTOH with this new logic
(attempt 2 times), in certain cases backends will get WALWriteLock,
when it is not really required which I am not sure can cause any problem,
so lets retain it as you have done in patch.

Post by Andres Freund

Post by Amit Kapila
11.
LWLockAcquireOrWait()
{
..
Assert(mode == LW_SHARED || mode == LW_EXCLUSIVE);
}
Isn't it better to use AssertArg() rather than Assert in above usage?

I've changed that. But I have to admit, I don't really see the point of
AssertArg(). If it'd output the value, then it'd be beneficial, but it
doesn't.

It seems you forgot to change in LWLockAcquireOrWait(), refer
below code.
@@ -773,14 +1191,14 @@ LWLockAcquireOrWait(LWLock *lock, LWLockMode mode)
int extraWaits = 0;
#ifdef LWLOCK_STATS
lwlock_stats *lwstats;
-#endif
-
- PRINT_LWDEBUG("LWLockAcquireOrWait", lock);

-#ifdef LWLOCK_STATS
lwstats = get_lwlock_stats_entry(lock);
#endif

+ Assert(mode == LW_SHARED || mode == LW_EXCLUSIVE);
+

Post by Andres Freund

Post by Amit Kapila
15.
/* must be greater than MAX_BACKENDS */
#define SHARED_LOCK_MASK (~EXCLUSIVE_LOCK)
a. how can we guarantee it to be greater than MaxBackends,
as MaxBackends is int (the max value for which will be
equal to SHARED_LOCK_MASK)?

MaxBackends luckily is limited to something lower. I've added a comment
to that regard.

Post by Amit Kapila
b. This is used only for LWLOCK_STATS, so shouldn't we
define this under LWLOCK_STATS.

It's a general value, so I don't think that's appropriate.

No issues.


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Today, I have verified all previous comments raised by

Post by Andres Freund

Post by Amit Kapila
4.
LWLockAcquireCommon()
{
..
if (!LWLockDequeueSelf(l))
{
for (;;)
{
PGSemaphoreLock(&proc->sem, false);
if (!proc->lwWaiting)
break;
extraWaits++;
}
lock->releaseOK = true;
..
}
Setting releaseOK in above context might not be required because if the
control comes in this part of code, it will not retry to acquire another
time.

Hm. You're probably right.

You have agreed to fix this comment, but it seems you have forgot
to change it.

Post by Andres Freund

Post by Amit Kapila
11.
LWLockRelease()
{
..
PRINT_LWDEBUG("LWLockRelease", lock, mode);
}
Shouldn't this be in begining of LWLockRelease function rather than
after processing held_lwlocks array?

Ok.

You have agreed to fix this comment, but it seems you have forgot
to change it.

Below comment doesn't seem to be adressed?

Post by Andres Freund
LWLockAcquireOrWait()
{
..
LOG_LWDEBUG("LWLockAcquireOrWait", lockid, mode, "succeeded");
..
}
a. such a log is not there in any other LWLock.. variants,
if we want to introduce it, then shouldn't it be done at
other places as well.

Below point is yet to be resolved.

Post by Andres Freund

Post by Amit Kapila
12.
#ifdef LWLOCK_DEBUG
lock->owner = MyProc;
#endif
Shouldn't it be reset in LWLockRelease?

That's actually intentional. It's quite useful to know the last owner
when debugging lwlock code.

Won't it cause any problem if the last owner process exits?


Can you explain how pg_read_barrier() in below code makes this
access safe?

LWLockWakeup()
{
..
+ pg_read_barrier(); /* pairs with nwaiters-- */
+ if (!BOOL_ACCESS_ONCE(lock->releaseOK))
..
}



With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
12.
#ifdef LWLOCK_DEBUG
lock->owner = MyProc;
#endif
Shouldn't it be reset in LWLockRelease?

That's actually intentional. It's quite useful to know the last owner
when debugging lwlock code.

Won't it cause any problem if the last owner process exits?

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Christian Kruse]

Hi,

I'm doing some benchmarks regarding this problem: one set with
baseline and one set with your patch. Machine was a 32 core machine (4
CPUs with 8 cores), 252 gib RAM. Both versions have the type align
patch applied. pgbench-tools config:

SCALES="100"
SETCLIENTS="1 4 8 16 32 48 64 96 128"
SETTIMES=2

I added -M prepared to the pgbench call in the benchwarmer script.

The read-only tests are finished, I come to similiar results as yours:

<http://wwwtech.de/pg/benchmarks-lwlock-read-only/>

I think the small differences are caused by the fact that I use TCP
connections and not Unix domain sockets.

The results are pretty impressive
 I will post the read-write results
as soon as they are finished.

Best regards,

[Peter Geoghegan]

On Tue, Feb 4, 2014 at 11:39 AM, Christian Kruse

Post by Christian Kruse
I'm doing some benchmarks regarding this problem: one set with
baseline and one set with your patch. Machine was a 32 core machine (4
CPUs with 8 cores), 252 gib RAM. Both versions have the type align
patch applied.

It certainly seems as if the interesting cases are where clients > cores.

[Peter Geoghegan]

On Tue, Feb 4, 2014 at 11:39 AM, Christian Kruse

Post by Christian Kruse
I added -M prepared to the pgbench call in the benchwarmer script.
<http://wwwtech.de/pg/benchmarks-lwlock-read-only/>

Note that Christian ran this test with max_connections=201, presumably
to exercise the alignment problem.

[Andres Freund]

Post by Peter Geoghegan
On Tue, Feb 4, 2014 at 11:39 AM, Christian Kruse

Post by Christian Kruse
I added -M prepared to the pgbench call in the benchwarmer script.
<http://wwwtech.de/pg/benchmarks-lwlock-read-only/>

Note that Christian ran this test with max_connections=201, presumably
to exercise the alignment problem.

I think he has applied the patch to hack around the alignment issue I
pushed to git for both branches. It's not nice enough to be applied yet,
but it should fix the issue.
I think the 201 is just a remembrance of debugging the issue.

Greetings,

Andres Freund

[Peter Geoghegan]

Post by Andres Freund
I think he has applied the patch to hack around the alignment issue I
pushed to git for both branches. It's not nice enough to be applied yet,
but it should fix the issue.
I think the 201 is just a remembrance of debugging the issue.

I guess that given that *both* cases tested had the patch applied,
that makes sense. However, I would have liked to see a real master
baseline.

[Andres Freund]

Post by Peter Geoghegan

Post by Andres Freund
I think he has applied the patch to hack around the alignment issue I
pushed to git for both branches. It's not nice enough to be applied

yet,

Post by Andres Freund
but it should fix the issue.
I think the 201 is just a remembrance of debugging the issue.

I guess that given that *both* cases tested had the patch applied,
that makes sense. However, I would have liked to see a real master
baseline.

Christian, could you rerun with master (the commit on which the branch is based on), the alignment patch, and then the lwlock patch? Best with max_connections 200.
That's probably more important than the write tests as a first step..

Thanks,
Andres
--
Please excuse brevity and formatting - I am writing this on my mobile phone.

Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Christian Kruse]

Hi,

Post by Andres Freund
Christian, could you rerun with master (the commit on which the
branch is based on), the alignment patch, and then the lwlock patch?
Best with max_connections 200. That's probably more important than
the write tests as a first step..

Ok, benchmark for baseline+alignment patch is running. This will take
a couple of hours and since I have to get up at about 05:00 I won't be
able to post it before tomorrow.

Best regards,

[Peter Geoghegan]

On Tue, Feb 4, 2014 at 12:30 PM, Christian Kruse

Post by Christian Kruse
Ok, benchmark for baseline+alignment patch is running.

I see that you have enabled latency information. For this kind of
thing I prefer to hack pgbench-tools to not collect this (i.e. to not
pass the "-l" flag, "Per-Transaction Logging"). Just remove it and
pgbench-tools rolls with it. It may well be that the overhead added is
completely insignificant, but for something like this, where the
latency information is unlikely to add any value, I prefer to not take
the chance. This is a fairly minor point, however, especially since
these are only 60 second runs where you're unlikely to accumulate
enough transaction latency information to notice any effect.

[Andres Freund]

Post by Peter Geoghegan
On Tue, Feb 4, 2014 at 12:30 PM, Christian Kruse

Post by Christian Kruse
Ok, benchmark for baseline+alignment patch is running.

I see that you have enabled latency information. For this kind of
thing I prefer to hack pgbench-tools to not collect this (i.e. to not
pass the "-l" flag, "Per-Transaction Logging"). Just remove it and
pgbench-tools rolls with it. It may well be that the overhead added is
completely insignificant, but for something like this, where the
latency information is unlikely to add any value, I prefer to not take
the chance. This is a fairly minor point, however, especially since
these are only 60 second runs where you're unlikely to accumulate
enough transaction latency information to notice any effect.

Hm, I don't find that convincing. If you look at the results from the
last run the latency information is actually quite interesting.

Greetings,

Andres Freund

[Peter Geoghegan]

On Tue, Feb 4, 2014 at 11:39 AM, Christian Kruse

Post by Christian Kruse
I'm doing some benchmarks regarding this problem: one set with
baseline and one set with your patch. Machine was a 32 core machine (4
CPUs with 8 cores), 252 gib RAM.

What CPU model? Can you post /proc/cpuinfo? The distinction between
logical and physical cores matters here.

[Christian Kruse]

Hi,

Post by Peter Geoghegan
On Tue, Feb 4, 2014 at 11:39 AM, Christian Kruse

Post by Christian Kruse
I'm doing some benchmarks regarding this problem: one set with
baseline and one set with your patch. Machine was a 32 core machine (4
CPUs with 8 cores), 252 gib RAM.

What CPU model? Can you post /proc/cpuinfo? The distinction between
logical and physical cores matters here.

model name : Intel(R) Xeon(R) CPU E5-4620 0 @ 2.20GHz

32 physical cores, 64 logical cores. /proc/cpuinfo is applied.

Best regards,

[Christian Kruse]

Hi,

[
] /proc/cpuinfo is applied.

I meant „attached“ and it seems that scp on /proc/cpuinfo doesn't
work


This time the cpuinfo is attached with content ;-)

Best regards,

[Andres Freund]

Hi,

Over at -performance Mark Kirkwood tested a recent version of this
(http://archives.postgresql.org/message-id/53B283F3.7020005%40catalyst.net.nz)

Test: pgbench
Options: scale 500
read only
Os: Ubuntu 14.04
Pg: 9.3.4
max_connections = 200
shared_buffers = 10GB
maintenance_work_mem = 1GB
effective_io_concurrency = 10
wal_buffers = 32MB
checkpoint_segments = 192
checkpoint_completion_target = 0.8
Results
Clients | 9.3 tps 32 cores | 9.3 tps 60 cores
--------+------------------+-----------------
6 | 70400 | 71028
12 | 98918 | 129140
24 | 230345 | 240631
48 | 324042 | 409510
96 | 346929 | 120464
192 | 312621 | 92663
So we have anti scaling with 60 cores as we increase the client connections.
Ouch! A level of urgency led to trying out Andres's 'rwlock' 9.4 branch [1]
- cherry picking the last 5 commits into 9.4 branch and building a package
Clients | 9.4 tps 60 cores (rwlock)
--------+--------------------------
6 | 70189
12 | 128894
24 | 233542
48 | 422754
96 | 590796
192 | 630672

Now, this is a bit of a skewed comparison due to 9.4 vs. 9.3 but still
interesting.

Greetings,

Andres Freund

[Heikki Linnakangas]

Post by Andres Freund
Hi,
Over at -performance Mark Kirkwood tested a recent version of this
(http://archives.postgresql.org/message-id/53B283F3.7020005%40catalyst.net.nz)

Test: pgbench
Options: scale 500
read only
Os: Ubuntu 14.04
Pg: 9.3.4
max_connections = 200
shared_buffers = 10GB
maintenance_work_mem = 1GB
effective_io_concurrency = 10
wal_buffers = 32MB
checkpoint_segments = 192
checkpoint_completion_target = 0.8
Results
Clients | 9.3 tps 32 cores | 9.3 tps 60 cores
--------+------------------+-----------------
6 | 70400 | 71028
12 | 98918 | 129140
24 | 230345 | 240631
48 | 324042 | 409510
96 | 346929 | 120464
192 | 312621 | 92663
So we have anti scaling with 60 cores as we increase the client connections.
Ouch! A level of urgency led to trying out Andres's 'rwlock' 9.4 branch [1]
- cherry picking the last 5 commits into 9.4 branch and building a package
Clients | 9.4 tps 60 cores (rwlock)
--------+--------------------------
6 | 70189
12 | 128894
24 | 233542
48 | 422754
96 | 590796
192 | 630672

Now, this is a bit of a skewed comparison due to 9.4 vs. 9.3 but still
interesting.

It looks like the issue I reported here:

http://www.postgresql.org/message-id/***@vmware.com

fixed by this commit:

http://git.postgresql.org/gitweb/?p=postgresql.git;a=commit;h=b03d196be055450c7260749f17347c2d066b4254.

So, definitely need to compare plain 9.4 vs patched 9.4, not 9.3.

- Heikki

[Mark Kirkwood]

Post by Heikki Linnakangas

Post by Andres Freund
Hi,
Over at -performance Mark Kirkwood tested a recent version of this
(http://archives.postgresql.org/message-id/53B283F3.7020005%40catalyst.net.nz)

Test: pgbench
Options: scale 500
read only
Os: Ubuntu 14.04
Pg: 9.3.4
max_connections = 200
shared_buffers = 10GB
maintenance_work_mem = 1GB
effective_io_concurrency = 10
wal_buffers = 32MB
checkpoint_segments = 192
checkpoint_completion_target = 0.8
Results
Clients | 9.3 tps 32 cores | 9.3 tps 60 cores
--------+------------------+-----------------
6 | 70400 | 71028
12 | 98918 | 129140
24 | 230345 | 240631
48 | 324042 | 409510
96 | 346929 | 120464
192 | 312621 | 92663
So we have anti scaling with 60 cores as we increase the client connections.
Ouch! A level of urgency led to trying out Andres's 'rwlock' 9.4 branch [1]
- cherry picking the last 5 commits into 9.4 branch and building a package
Clients | 9.4 tps 60 cores (rwlock)
--------+--------------------------
6 | 70189
12 | 128894
24 | 233542
48 | 422754
96 | 590796
192 | 630672

Now, this is a bit of a skewed comparison due to 9.4 vs. 9.3 but still
interesting.

http://git.postgresql.org/gitweb/?p=postgresql.git;a=commit;h=b03d196be055450c7260749f17347c2d066b4254.
So, definitely need to compare plain 9.4 vs patched 9.4, not 9.3.

Here's plain 9.4 vs patched 9.4:

Clients | 9.4 tps 60 cores | 9.4 tps 60 cores (rwlock)
--------+------------------+--------------------------
6 | 69490 | 70189
12 | 128200 | 128894
24 | 232243 | 233542
48 | 417689 | 422754
96 | 464037 | 590796
192 | 418252 | 630672

It appears that plain 9.4 does not exhibit the dramatic anti scaling
that 9.3 showed, but there is still evidence of some contention in the
higher client numbers, and we peak at the 96 client mark. The patched
variant looks pretty much free from this, still scaling at 192
connections (might have been interesting to try more, but had
max_connections set to 200)!

Cheers

Mark

[Andres Freund]

Hi,

Attached you can find the next version of my LW_SHARED patchset. Now
that atomics are committed, it seems like a good idea to also add their
raison d'être.

Since the last public version I have:
* Addressed lots of Amit's comments. Thanks!
* Peformed a fair amount of testing.
* Rebased the code. The volatile removal made that not entirely
trivial...
* Significantly cleaned up and simplified the code.
* Updated comments and such
* Fixed a minor bug (unpaired HOLD/RESUME_INTERRUPTS in a corner case)

The feature currently consists out of two patches:
1) Convert PGPROC->lwWaitLink into a dlist. The old code was frail and
verbose. This also does:
* changes the logic in LWLockRelease() to release all shared lockers
when waking up any. This can yield some significant performance
improvements - and the fairness isn't really much worse than
before,
as we always allowed new shared lockers to jump the queue.

* adds a memory pg_write_barrier() in the wakeup paths between
dequeuing and unsetting ->lwWaiting. That was always required on
weakly ordered machines, but f4077cda2 made it more urgent. I can
reproduce crashes without it.
2) Implement the wait free LW_SHARED algorithm.

Personally I'm quite happy with the new state. I think it needs more
review, but I personally don't know of anything that needs
changing. There's lots of further improvements that could be done, but
let's get this in first.

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Robert Haas]

Post by Andres Freund
1) Convert PGPROC->lwWaitLink into a dlist. The old code was frail and
* changes the logic in LWLockRelease() to release all shared lockers
when waking up any. This can yield some significant performance
improvements - and the fairness isn't really much worse than
before,
as we always allowed new shared lockers to jump the queue.
* adds a memory pg_write_barrier() in the wakeup paths between
dequeuing and unsetting ->lwWaiting. That was always required on
weakly ordered machines, but f4077cda2 made it more urgent. I can
reproduce crashes without it.

I think it's a really bad idea to mix a refactoring change (like
converting PGPROC->lwWaitLink into a dlist) with an attempted
performance enhancement (like changing the rules for jumping the lock
queue) and a bug fix (like adding pg_write_barrier where needed). I'd
suggest that the last of those be done first, and perhaps
back-patched.

The current coding, using a hand-rolled list, touches shared memory
fewer times. When many waiters are awoken at once, we clip them all
out of the list at one go. Your revision moves them to a
backend-private list one at a time, and then pops them off one at a
time. The backend-private memory accesses don't seem like they matter
much, but the shared memory accesses would be nice to avoid.

Does LWLockUpdateVar's wake-up loop need a write barrier per
iteration, or just one before the loop starts? How about commenting
the pg_write_barrier() with the read-fence to which it pairs?

+ if(waiter->lwWaitMode == LW_EXCLUSIVE)

Whitespace.
--
Robert Haas
EnterpriseDB: http://www.enterprisedb.com
The Enterprise PostgreSQL Company

[Andres Freund]

Post by Robert Haas

Post by Andres Freund
1) Convert PGPROC->lwWaitLink into a dlist. The old code was frail and
* changes the logic in LWLockRelease() to release all shared lockers
when waking up any. This can yield some significant performance
improvements - and the fairness isn't really much worse than
before,
as we always allowed new shared lockers to jump the queue.
* adds a memory pg_write_barrier() in the wakeup paths between
dequeuing and unsetting ->lwWaiting. That was always required on
weakly ordered machines, but f4077cda2 made it more urgent. I can
reproduce crashes without it.

I think it's a really bad idea to mix a refactoring change (like
converting PGPROC->lwWaitLink into a dlist) with an attempted
performance enhancement (like changing the rules for jumping the lock
queue) and a bug fix (like adding pg_write_barrier where needed). I'd
suggest that the last of those be done first, and perhaps
back-patched.

I think it makes sense to separate out the write barrier one. I don't
really see the point of separating the other two changes.

I've indeed previously started a thread
(http://archives.postgresql.org/message-id/20140210134625.GA15246%40awork2.anarazel.de)
about the barrier issue. IIRC you argued that that might be to
expensive.

Post by Robert Haas
The current coding, using a hand-rolled list, touches shared memory
fewer times. When many waiters are awoken at once, we clip them all
out of the list at one go. Your revision moves them to a
backend-private list one at a time, and then pops them off one at a
time. The backend-private memory accesses don't seem like they matter
much, but the shared memory accesses would be nice to avoid.

I can't imagine this to matter. We're entering the kernel for each PROC
for the PGSemaphoreUnlock() and we're dirtying the cacheline for
proc->lwWaiting = false anyway. This really is the slow path.

Post by Robert Haas
Does LWLockUpdateVar's wake-up loop need a write barrier per
iteration, or just one before the loop starts? How about commenting
the pg_write_barrier() with the read-fence to which it pairs?

Hm. Are you picking out LWLockUpdateVar for a reason or just as an
example? Because I don't see a difference between the different wakeup
loops?
It needs to be a barrier per iteration.

Currently the loop looks like
while (head != NULL)
{
proc = head;
head = proc->lwWaitLink;
proc->lwWaitLink = NULL;
proc->lwWaiting = false;
PGSemaphoreUnlock(&proc->sem);
}

Consider what happens if either the compiler or the cpu reorders this
to:
proc->lwWaiting = false;
head = proc->lwWaitLink;
proc->lwWaitLink = NULL;
PGSemaphoreUnlock(&proc->sem);

as soon as lwWaiting = false, 'proc' can wake up and acquire a new
lock. Backends can wake up prematurely because proc->sem is used for
other purposes than this (that's why the loops around PGSemaphoreLock
exist). Then it could reset lwWaitLink while acquiring a new lock. And
some processes wouldn't be woken up anymore.

The barrier it pairs with is the spinlock acquiration before
requeuing. To be more obviously correct we could add a read barrier
before
if (!proc->lwWaiting)
break;
but I don't think it's needed.

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Robert Haas]

Post by Andres Freund
2) Implement the wait free LW_SHARED algorithm.

+ * too high for workloads/locks that were locked in shared mode very

s/locked/taken/?

+ * frequently. Often we were spinning in the (obviously exlusive) spinlock,

exclusive.

+ * acquiration for locks that aren't exclusively locked.

acquisition.

+ * For exlusive lock acquisition we use an atomic compare-and-exchange on the

exclusive.

+ * lockcount variable swapping in EXCLUSIVE_LOCK/1<<31-1/0x7FFFFFFF if and only

Add comma after variable. Find some way of describing the special
value (maybe "a sentinel value, EXCLUSIVE_LOCK") just once, instead of
three times.

+ * if the current value of lockcount is 0. If the swap was not successfull, we

successful.

+ * by 1 again. If so, we have to wait for the exlusive locker to release the

exclusive.

+ * The attentive reader probably might have noticed that naively doing the

"probably might" is redundant. Delete probably.

+ * notice that we have to wait. Unfortunately until we have finished queuing,

until -> by the time

+ * Phase 2: Add us too the waitqueue of the lock

too -> to. And maybe us -> ourselves.

+ * get queued in Phase 2 and we can wake them up if neccessary or they will

necessary.

+ * When acquiring shared locks it's possible that we disturb an exclusive
+ * waiter. If that's a problem for the specific user, pass in a valid pointer
+ * for 'potentially_spurious'. Its value will be set to true if we possibly
+ * did so. The caller then has to handle that scenario.

"disturb" is not clear enough.

+ /* yipeyyahee */

Although this will be clear to individuals with a good command of
English, I suggest avoiding such usages.
--
Robert Haas
EnterpriseDB: http://www.enterprisedb.com
The Enterprise PostgreSQL Company

[Andres Freund]

Hi Robert,

[ comment fixes ]

Thanks, I've incorporated these + a bit more.

Could you otherwise make sense of the explanation and the algorithm?

+ /* yipeyyahee */
Although this will be clear to individuals with a good command of
English, I suggest avoiding such usages.

I've removed them with a heavy heart. These are heartfelt emotions from
getting the algorithm to work.... (:P)

I've attached these fixes + the removal of spinlocks around releaseOK as
follow up patches. Obviously they'll be merged into the other patch, but
sounds useful to be able see them separately.

Greetings,

Andres Freund

[Jim Nasby]

+#define EXCLUSIVE_LOCK (((uint32) 1) << (31 - 1))
+
+/* Must be greater than MAX_BACKENDS - which is 2^23-1, so we're fine. */
+#define SHARED_LOCK_MASK (~EXCLUSIVE_LOCK)

There should at least be a comment where we define MAX_BACKENDS about the relationship here... or better yet, validate that MAX_BACKENDS > SHARED_LOCK_MASK during postmaster startup. (For those that think that's too pedantic, I'll argue that it's no worse than the patch verifying that MyProc != NULL in LWLockQueueSelf()).

+/*
+ * Internal function that tries to atomically acquire the lwlock in the passed
+ * in mode.
+ *
+ * This function will not block waiting for a lock to become free - that's the
+ * callers job.
+ *
+ * Returns true if the lock isn't free and we need to wait.
+ *
+ * When acquiring shared locks it's possible that we disturb an exclusive
+ * waiter. If that's a problem for the specific user, pass in a valid pointer
+ * for 'potentially_spurious'. Its value will be set to true if we possibly
+ * did so. The caller then has to handle that scenario.
+ */
+static bool
+LWLockAttemptLock(LWLock* lock, LWLockMode mode, bool *potentially_spurious)

We should invert the return of this function. Current code returns true if the lock is actually acquired (see below), and I think that's true of other locking code as well. IMHO it makes more sense that way, plus consistency is good.

(From 9.3)
* LWLockConditionalAcquire - acquire a lightweight lock in the specified mode
*
* If the lock is not available, return FALSE with no side-effects.

[Andres Freund]

Post by Jim Nasby

+#define EXCLUSIVE_LOCK (((uint32) 1) << (31 - 1))
+
+/* Must be greater than MAX_BACKENDS - which is 2^23-1, so we're fine. */
+#define SHARED_LOCK_MASK (~EXCLUSIVE_LOCK)

There should at least be a comment where we define MAX_BACKENDS about the relationship here... or better yet, validate that MAX_BACKENDS > SHARED_LOCK_MASK during postmaster startup. (For those that think that's too pedantic, I'll argue that it's no worse than the patch verifying that MyProc != NULL in LWLockQueueSelf()).

If you modify either, you better grep for them... I don't think that's
going to happen anyway. Requiring it during startup would mean exposing
SHARED_LOCK_MASK outside of lwlock.c which'd be ugly. We could possibly
stick a StaticAssert() someplace in lwlock.c.

And no, it's not comparable at all to MyProc != NULL - the lwlock code
initially *does* run when MyProc isn't setup. We just better not
conflict against any other lockers at that stage.

Post by Jim Nasby

+/*
+ * Internal function that tries to atomically acquire the lwlock in the passed
+ * in mode.
+ *
+ * This function will not block waiting for a lock to become free - that's the
+ * callers job.
+ *
+ * Returns true if the lock isn't free and we need to wait.
+ *
+ * When acquiring shared locks it's possible that we disturb an exclusive
+ * waiter. If that's a problem for the specific user, pass in a valid pointer
+ * for 'potentially_spurious'. Its value will be set to true if we possibly
+ * did so. The caller then has to handle that scenario.
+ */
+static bool
+LWLockAttemptLock(LWLock* lock, LWLockMode mode, bool *potentially_spurious)

We should invert the return of this function. Current code returns
true if the lock is actually acquired (see below), and I think that's
true of other locking code as well. IMHO it makes more sense that way,
plus consistency is good.

I don't think so. I've wondered about it as well, but the way the
function is used its more consistent imo if it returns whether we must
wait. Note that it's not an exported function.

Greetings,

Andres Freund

[Jim Nasby]

Post by Andres Freund
If you modify either, you better grep for them... I don't think that's
going to happen anyway. Requiring it during startup would mean exposing
SHARED_LOCK_MASK outside of lwlock.c which'd be ugly. We could possibly
stick a StaticAssert() someplace in lwlock.c.

Ahh, yeah, exposing it would be ugly.

I just get the heeby-jeebies when I see assumptions like this though. I fear there's a bunch of cases where changing something will break a completely unrelated part of the system with no warning.

Maybe add an assert() to check it?

Post by Andres Freund
And no, it's not comparable at all to MyProc != NULL - the lwlock code
initially*does* run when MyProc isn't setup. We just better not
conflict against any other lockers at that stage.

Ahh, can you maybe add that detail to the comment? That wasn't clear to me.

Post by Andres Freund

Post by Jim Nasby

+/*
+ * Internal function that tries to atomically acquire the lwlock in the passed
+ * in mode.
+ *
+ * This function will not block waiting for a lock to become free - that's the
+ * callers job.
+ *
+ * Returns true if the lock isn't free and we need to wait.
+ *
+ * When acquiring shared locks it's possible that we disturb an exclusive
+ * waiter. If that's a problem for the specific user, pass in a valid pointer
+ * for 'potentially_spurious'. Its value will be set to true if we possibly
+ * did so. The caller then has to handle that scenario.
+ */
+static bool
+LWLockAttemptLock(LWLock* lock, LWLockMode mode, bool *potentially_spurious)

We should invert the return of this function. Current code returns
true if the lock is actually acquired (see below), and I think that's
true of other locking code as well. IMHO it makes more sense that way,
plus consistency is good.

I don't think so. I've wondered about it as well, but the way the
function is used its more consistent imo if it returns whether we must
wait. Note that it's not an exported function.

ISTM that a function attempting a lock would return success, not failure. Even though it's internal now it could certainly be made external at some point in the future. But I suppose it's ultimately a matter of preference...

[Amit Kapila]

Post by Andres Freund
Hi,
Attached you can find the next version of my LW_SHARED patchset. Now
that atomics are committed, it seems like a good idea to also add their
raison d'être.
* Addressed lots of Amit's comments. Thanks!
* Peformed a fair amount of testing.
* Rebased the code. The volatile removal made that not entirely
trivial...
* Significantly cleaned up and simplified the code.
* Updated comments and such
* Fixed a minor bug (unpaired HOLD/RESUME_INTERRUPTS in a corner case)
1) Convert PGPROC->lwWaitLink into a dlist. The old code was frail and
* changes the logic in LWLockRelease() to release all shared lockers
when waking up any. This can yield some significant performance
improvements - and the fairness isn't really much worse than
before,
as we always allowed new shared lockers to jump the queue.
* adds a memory pg_write_barrier() in the wakeup paths between
dequeuing and unsetting ->lwWaiting. That was always required on
weakly ordered machines, but f4077cda2 made it more urgent. I can
reproduce crashes without it.
2) Implement the wait free LW_SHARED algorithm.

I have done few performance tests for above patches and results of
same is as below:

Performance Data
------------------------------
IBM POWER-7 16 cores, 64 hardware threads
RAM = 64GB
max_connections =210
Database Locale =C
checkpoint_segments=256
checkpoint_timeout =35min
shared_buffers=8GB
Client Count = number of concurrent sessions and threads (ex. -c 8 -j 8)
Duration of each individual run = 5mins
Test type - read only pgbench with -M prepared
Other Related information about test
a. This is the data for median of 3 runs, the detailed data of individual
run
is attached with mail.
b. I have applied both the patches to take performance data.

Scale Factor - 100

Patch_ver/Client_count 1 8 16 32 64 128 HEAD 13344 106921 196629 295123
377846 333928 PATCH 13662 106179 203960 298955 452638 465671

Scale Factor - 3000

Patch_ver/Client_count 8 16 32 64 128 160 HEAD 86920 152417 231668
280827 257093 255122 PATCH 87552 160313 230677 276186 248609 244372


Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all the
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client count,
however after that it starts dipping (in above config unto ~4.4%).

The above data shows that the patch improves performance for cases
when there is shared LWLock contention, however there is a slight
performance dip in case of Exclusive LWLocks (at scale factor 3000,
it needs exclusive LWLocks for buf mapping tables). Now I am not
sure if this is the worst case dip or under similar configurations the
performance dip can be higher, because the trend shows that dip is
increasing with more client counts.

Brief Analysis of code w.r.t performance dip
---------------------------------------------------------------------
Extra Instructions w.r.t Head in Acquire Exclusive lock path
a. Attempt lock twice
b. atomic operations for nwaiters in LWLockQueueSelf() and
LWLockAcquireCommon()
c. Now we need to take spinlock twice, once for self queuing and then
again for setting releaseOK.
d. few function calls and some extra checks

Similarly there seems to be few additional instructions in
LWLockRelease() path.

Now probably these shouldn't matter much in case backend needs to
wait for other Exclusive locker, but I am not sure what else could be
the reason for dip in case we need to have Exclusive LWLocks.


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Hi,

Post by Amit Kapila
I have done few performance tests for above patches and results of

Cool, thanks.

Post by Amit Kapila
Performance Data
------------------------------
IBM POWER-7 16 cores, 64 hardware threads
RAM = 64GB
max_connections =210
Database Locale =C
checkpoint_segments=256
checkpoint_timeout =35min
shared_buffers=8GB
Client Count = number of concurrent sessions and threads (ex. -c 8 -j 8)
Duration of each individual run = 5mins
Test type - read only pgbench with -M prepared
Other Related information about test
a. This is the data for median of 3 runs, the detailed data of individual
run
is attached with mail.
b. I have applied both the patches to take performance data.
Scale Factor - 100
Patch_ver/Client_count 1 8 16 32 64 128 HEAD 13344 106921 196629 295123
377846 333928 PATCH 13662 106179 203960 298955 452638 465671
Scale Factor - 3000
Patch_ver/Client_count 8 16 32 64 128 160 HEAD 86920 152417 231668
280827 257093 255122 PATCH 87552 160313 230677 276186 248609 244372
Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all the
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client count,
however after that it starts dipping (in above config unto ~4.4%).

Hm. Interesting. I don't see that dip on x86.

Post by Amit Kapila
The above data shows that the patch improves performance for cases
when there is shared LWLock contention, however there is a slight
performance dip in case of Exclusive LWLocks (at scale factor 3000,
it needs exclusive LWLocks for buf mapping tables). Now I am not
sure if this is the worst case dip or under similar configurations the
performance dip can be higher, because the trend shows that dip is
increasing with more client counts.
Brief Analysis of code w.r.t performance dip
---------------------------------------------------------------------
Extra Instructions w.r.t Head in Acquire Exclusive lock path
a. Attempt lock twice
b. atomic operations for nwaiters in LWLockQueueSelf() and
LWLockAcquireCommon()
c. Now we need to take spinlock twice, once for self queuing and then
again for setting releaseOK.
d. few function calls and some extra checks

Hm. I can't really see the number of atomics itself matter - a spinning
lock will do many more atomic ops than this. But I wonder whether we
could get rid of the releaseOK lock. Should be quite possible.

Post by Amit Kapila
Now probably these shouldn't matter much in case backend needs to
wait for other Exclusive locker, but I am not sure what else could be
the reason for dip in case we need to have Exclusive LWLocks.

Any chance to get a profile?

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
I have done few performance tests for above patches and results of

Cool, thanks.

Post by Amit Kapila
Performance Data
------------------------------
IBM POWER-7 16 cores, 64 hardware threads
RAM = 64GB
max_connections =210
Database Locale =C
checkpoint_segments=256
checkpoint_timeout =35min
shared_buffers=8GB
Client Count = number of concurrent sessions and threads (ex. -c 8 -j 8)
Duration of each individual run = 5mins
Test type - read only pgbench with -M prepared
Other Related information about test
a. This is the data for median of 3 runs, the detailed data of individual
run
is attached with mail.
b. I have applied both the patches to take performance data.
Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all the
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client count,
however after that it starts dipping (in above config unto ~4.4%).

Hm. Interesting. I don't see that dip on x86.

Is it possible that implementation of some atomic operation is costlier
for particular architecture?

I have tried again for scale factor 3000 and could see the dip and this
time I have even tried with 175 client count and the dip is approximately
5% which is slightly more than 160 client count.


Patch_ver/Client_count 175 HEAD 248374 PATCH 235669

Post by Andres Freund

Post by Amit Kapila
Now probably these shouldn't matter much in case backend needs to
wait for other Exclusive locker, but I am not sure what else could be
the reason for dip in case we need to have Exclusive LWLocks.

Any chance to get a profile?

Here it goes..

HEAD - client_count=128
-----------------------------------------

+ 7.53% postgres postgres [.] GetSnapshotData
+ 3.41% postgres postgres [.] AllocSetAlloc
+ 2.61% postgres postgres [.] AllocSetFreeIndex
+ 2.49% postgres postgres [.] _bt_compare
+ 2.43% postgres [kernel.kallsyms] [k] .__copy_tofrom_user
+ 2.40% postgres postgres [.]
hash_search_with_hash_value
+ 1.83% postgres postgres [.] tas
+ 1.29% postgres postgres [.] pg_encoding_mbcliplen
+ 1.27% postgres postgres [.] MemoryContextCreate
+ 1.22% postgres postgres [.]
MemoryContextAllocZeroAligned
+ 1.17% postgres postgres [.] hash_seq_search
+ 0.97% postgres postgres [.] LWLockRelease
+ 0.96% postgres postgres [.]
MemoryContextAllocZero
+ 0.91% postgres postgres [.]
GetPrivateRefCountEntry
+ 0.82% postgres postgres [.] AllocSetFree
+ 0.79% postgres postgres [.] LWLockAcquireCommon
+ 0.78% postgres postgres [.] pfree



Detailed Data
-------------------------
- 7.53% postgres postgres [.] GetSnapshotData
- GetSnapshotData
- 7.46% GetSnapshotData
- 7.46% GetTransactionSnapshot
- 3.74% exec_bind_message
PostgresMain
BackendRun
BackendStartup
ServerLoop
PostmasterMain
main
generic_start_main.isra.0
__libc_start_main
0
- 3.72% PortalStart
exec_bind_message
PostgresMain
BackendRun
BackendStartup
ServerLoop
PostmasterMain
main
generic_start_main.isra.0
__libc_start_main
0
- 3.41% postgres postgres [.] AllocSetAlloc
- AllocSetAlloc
- 2.01% AllocSetAlloc
0.81% palloc
0.63% MemoryContextAlloc
- 2.61% postgres postgres [.] AllocSetFreeIndex
- AllocSetFreeIndex
1.59% AllocSetAlloc
0.79% AllocSetFree
- 2.49% postgres postgres [.] _bt_compare
- _bt_compare
- 1.80% _bt_binsrch
- 1.80% _bt_binsrch
- 1.21% _bt_search
_bt_first

Lwlock_contention patches - client_count=128
----------------------------------------------------------------------

+ 7.95% postgres postgres [.] GetSnapshotData
+ 3.58% postgres postgres [.] AllocSetAlloc
+ 2.51% postgres postgres [.] _bt_compare
+ 2.44% postgres postgres [.]
hash_search_with_hash_value
+ 2.33% postgres [kernel.kallsyms] [k] .__copy_tofrom_user
+ 2.24% postgres postgres [.] AllocSetFreeIndex
+ 1.75% postgres postgres [.]
pg_atomic_fetch_add_u32_impl
+ 1.60% postgres postgres [.] hash_seq_search
+ 1.31% postgres postgres [.] pg_encoding_mbcliplen
+ 1.27% postgres postgres [.]
MemoryContextAllocZeroAligned
+ 1.26% postgres postgres [.] MemoryContextCreate
+ 0.98% postgres postgres [.] GetPrivateRefCountEntry
+ 0.97% postgres postgres [.] MemoryContextAllocZero
+ 0.87% postgres postgres [.] LWLockRelease
+ 0.82% postgres postgres [.] AllocSetFree
+ 0.79% postgres postgres [.] SearchCatCache
+ 0.70% postgres postgres [.] palloc
0.69% postgres postgres [.] pfree
+ 0.61% postgres postgres [.] AllocSetDelete
+ 0.57% postgres postgres [.] hash_any
0.57% postgres postgres [.] MemoryContextAlloc
0.56% postgres postgres [.] FunctionCall2Coll
+ 0.56% swapper [kernel.kallsyms] [k]
.pseries_dedicated_idle_sleep
0.56% postgres libc-2.14.90.so [.] memcpy
+ 0.55% postgres postgres [.] AllocSetReset
0.51% postgres postgres [.] _bt_binsrch
0.50% postgres postgres [.] LWLockAcquireCommon


Detailed Data
----------------------------
- 7.95% postgres postgres [.] GetSnapshotData

- GetSnapshotData


- 7.87% GetSnapshotData


- 7.87% GetTransactionSnapshot


- 3.95% exec_bind_message


0


- 3.92% PortalStart


exec_bind_message


0


- 3.58% postgres postgres [.] AllocSetAlloc


- AllocSetAlloc


- 1.82% AllocSetAlloc


0.75% palloc


0.56% MemoryContextAlloc


- 2.51% postgres postgres [.] _bt_compare


- _bt_compare


- 1.75% _bt_binsrch


- 1.75% _bt_binsrch


- 1.18% _bt_search


_bt_first


btgettuple


- 2.44% postgres postgres [.]
hash_search_with_hash_value

- hash_search_with_hash_value


- 2.01% hash_search_with_hash_value


- 0.96% BufTableLookup


BufferAlloc


- 2.33% postgres [kernel.kallsyms] [k] .__copy_tofrom_user


- .__copy_tofrom_user


- 2.27% .file_read_actor


.generic_file_aio_read


.do_sync_read


- 2.24% postgres postgres [.] AllocSetFreeIndex


- AllocSetFreeIndex


1.23% AllocSetAlloc


0.68% AllocSetFree


- 1.75% postgres postgres [.]
pg_atomic_fetch_add_u32_impl

- pg_atomic_fetch_add_u32_impl


0.95% pg_atomic_fetch_add_u32


0.75% pg_atomic_fetch_sub_u32_impl


- 1.60% postgres postgres [.] hash_seq_search


- hash_seq_search


- 0.91% PreCommit_Portals


- 0.91% PreCommit_Portals





With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Andres Freund

Post by Andres Freund

Post by Amit Kapila
Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all the
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client

count,

Post by Andres Freund

Post by Amit Kapila
however after that it starts dipping (in above config unto ~4.4%).

Hm. Interesting. I don't see that dip on x86.

Is it possible that implementation of some atomic operation is costlier
for particular architecture?

Yes, sure. And IIRC POWER improved atomics performance considerably for
POWER8...

Post by Andres Freund
I have tried again for scale factor 3000 and could see the dip and this
time I have even tried with 175 client count and the dip is approximately
5% which is slightly more than 160 client count.

FWIW, the profile always looks like
- 48.61% postgres postgres [.] s_lock
- s_lock
+ 96.67% StrategyGetBuffer
+ 1.19% UnpinBuffer
+ 0.90% PinBuffer
+ 0.70% hash_search_with_hash_value
+ 3.11% postgres postgres [.] GetSnapshotData
+ 2.47% postgres postgres [.] StrategyGetBuffer
+ 1.93% postgres [kernel.kallsyms] [k] copy_user_generic_string
+ 1.28% postgres postgres [.] hash_search_with_hash_value
- 1.27% postgres postgres [.] LWLockAttemptLock
- LWLockAttemptLock
- 97.78% LWLockAcquire
+ 38.76% ReadBuffer_common
+ 28.62% _bt_getbuf
+ 8.59% _bt_relandgetbuf
+ 6.25% GetSnapshotData
+ 5.93% VirtualXactLockTableInsert
+ 3.95% VirtualXactLockTableCleanup
+ 2.35% index_fetch_heap
+ 1.66% StartBufferIO
+ 1.56% LockReleaseAll
+ 1.55% _bt_next
+ 0.78% LockAcquireExtended
+ 1.47% _bt_next
+ 0.75% _bt_relandgetbuf

to me. Now that's with the client count 496, but it's similar with lower
counts.

BTW, that profile *clearly* indicates we should make StrategyGetBuffer()
smarter.

Post by Andres Freund
Patch_ver/Client_count 175 HEAD 248374 PATCH 235669

Post by Andres Freund

Post by Amit Kapila
Now probably these shouldn't matter much in case backend needs to
wait for other Exclusive locker, but I am not sure what else could be
the reason for dip in case we need to have Exclusive LWLocks.

Any chance to get a profile?

Here it goes..
Lwlock_contention patches - client_count=128
----------------------------------------------------------------------
+ 7.95% postgres postgres [.] GetSnapshotData
+ 3.58% postgres postgres [.] AllocSetAlloc
+ 2.51% postgres postgres [.] _bt_compare
+ 2.44% postgres postgres [.]
hash_search_with_hash_value
+ 2.33% postgres [kernel.kallsyms] [k] .__copy_tofrom_user
+ 2.24% postgres postgres [.] AllocSetFreeIndex
+ 1.75% postgres postgres [.]
pg_atomic_fetch_add_u32_impl

Uh. Huh? Normally that'll be inline. That's compiled with gcc? What were
the compiler settings you used?

Greetings,

Andres Freund

--
Andres Freund http://www.2ndQuadrant.com/
PostgreSQL Development, 24x7 Support, Training & Services

[Andres Freund]

Post by Andres Freund
FWIW, the profile always looks like
- 48.61% postgres postgres [.] s_lock
- s_lock
+ 96.67% StrategyGetBuffer
+ 1.19% UnpinBuffer
+ 0.90% PinBuffer
+ 0.70% hash_search_with_hash_value
+ 3.11% postgres postgres [.] GetSnapshotData
+ 2.47% postgres postgres [.] StrategyGetBuffer
+ 1.93% postgres [kernel.kallsyms] [k] copy_user_generic_string
+ 1.28% postgres postgres [.] hash_search_with_hash_value
- 1.27% postgres postgres [.] LWLockAttemptLock
- LWLockAttemptLock
- 97.78% LWLockAcquire
+ 38.76% ReadBuffer_common
+ 28.62% _bt_getbuf
+ 8.59% _bt_relandgetbuf
+ 6.25% GetSnapshotData
+ 5.93% VirtualXactLockTableInsert
+ 3.95% VirtualXactLockTableCleanup
+ 2.35% index_fetch_heap
+ 1.66% StartBufferIO
+ 1.56% LockReleaseAll
+ 1.55% _bt_next
+ 0.78% LockAcquireExtended
+ 1.47% _bt_next
+ 0.75% _bt_relandgetbuf
to me. Now that's with the client count 496, but it's similar with lower
counts.
BTW, that profile *clearly* indicates we should make StrategyGetBuffer()
smarter.

Which is nearly trivial now that atomics are in. Check out the attached
WIP patch which eliminates the spinlock from StrategyGetBuffer() unless
there's buffers on the freelist.

Test:
pgbench -M prepared -P 5 -S -c 496 -j 496 -T 5000
on a scale=1000 database, with 4GB of shared buffers.

Before:
progress: 40.0 s, 136252.3 tps, lat 3.628 ms stddev 4.547
progress: 45.0 s, 135049.0 tps, lat 3.660 ms stddev 4.515
progress: 50.0 s, 135788.9 tps, lat 3.640 ms stddev 4.398
progress: 55.0 s, 135268.4 tps, lat 3.654 ms stddev 4.469
progress: 60.0 s, 134991.6 tps, lat 3.661 ms stddev 4.739

after:
progress: 40.0 s, 207701.1 tps, lat 2.382 ms stddev 3.018
progress: 45.0 s, 208022.4 tps, lat 2.377 ms stddev 2.902
progress: 50.0 s, 209187.1 tps, lat 2.364 ms stddev 2.970
progress: 55.0 s, 206462.7 tps, lat 2.396 ms stddev 2.871
progress: 60.0 s, 210263.8 tps, lat 2.351 ms stddev 2.914

Yes, no kidding.

The results are similar, but less extreme, for smaller client counts
like 80 or 160.

Amit, since your test seems to be currently completely bottlenecked
within StrategyGetBuffer(), could you compare with that patch applied to
HEAD and the LW_SHARED patch for one client count? That'll may allow us
to see a more meaningful profile...

Greetings,

Andres Freund

[Merlin Moncure]

Post by Andres Freund

Post by Andres Freund
FWIW, the profile always looks like
- 48.61% postgres postgres [.] s_lock
- s_lock
+ 96.67% StrategyGetBuffer
+ 1.19% UnpinBuffer
+ 0.90% PinBuffer
+ 0.70% hash_search_with_hash_value
+ 3.11% postgres postgres [.] GetSnapshotData
+ 2.47% postgres postgres [.] StrategyGetBuffer
+ 1.93% postgres [kernel.kallsyms] [k] copy_user_generic_string
+ 1.28% postgres postgres [.] hash_search_with_hash_value
- 1.27% postgres postgres [.] LWLockAttemptLock
- LWLockAttemptLock
- 97.78% LWLockAcquire
+ 38.76% ReadBuffer_common
+ 28.62% _bt_getbuf
+ 8.59% _bt_relandgetbuf
+ 6.25% GetSnapshotData
+ 5.93% VirtualXactLockTableInsert
+ 3.95% VirtualXactLockTableCleanup
+ 2.35% index_fetch_heap
+ 1.66% StartBufferIO
+ 1.56% LockReleaseAll
+ 1.55% _bt_next
+ 0.78% LockAcquireExtended
+ 1.47% _bt_next
+ 0.75% _bt_relandgetbuf
to me. Now that's with the client count 496, but it's similar with lower
counts.
BTW, that profile *clearly* indicates we should make StrategyGetBuffer()
smarter.

Which is nearly trivial now that atomics are in. Check out the attached
WIP patch which eliminates the spinlock from StrategyGetBuffer() unless
there's buffers on the freelist.

Is this safe?

+ victim = pg_atomic_fetch_add_u32(&StrategyControl->nextVictimBuffer, 1);

- if (++StrategyControl->nextVictimBuffer >= NBuffers)
+ buf = &BufferDescriptors[victim % NBuffers];
+
+ if (victim % NBuffers == 0)
<snip>

I don't think there's any guarantee you could keep nextVictimBuffer
from wandering off the end. You could buy a little safety by CAS'ing
zero in instead, followed by atomic increment. However that's still
pretty dodgy IMO and requires two atomic ops which will underperform
the spin in some situations.

I like Robert's idea to keep the spinlock but preallocate a small
amount of buffers, say 8 or 16.

merlin

[Andres Freund]

Post by Merlin Moncure

Post by Andres Freund

Post by Andres Freund
FWIW, the profile always looks like
- 48.61% postgres postgres [.] s_lock
- s_lock
+ 96.67% StrategyGetBuffer
+ 1.19% UnpinBuffer
+ 0.90% PinBuffer
+ 0.70% hash_search_with_hash_value
+ 3.11% postgres postgres [.] GetSnapshotData
+ 2.47% postgres postgres [.] StrategyGetBuffer
+ 1.93% postgres [kernel.kallsyms] [k] copy_user_generic_string
+ 1.28% postgres postgres [.] hash_search_with_hash_value
- 1.27% postgres postgres [.] LWLockAttemptLock
- LWLockAttemptLock
- 97.78% LWLockAcquire
+ 38.76% ReadBuffer_common
+ 28.62% _bt_getbuf
+ 8.59% _bt_relandgetbuf
+ 6.25% GetSnapshotData
+ 5.93% VirtualXactLockTableInsert
+ 3.95% VirtualXactLockTableCleanup
+ 2.35% index_fetch_heap
+ 1.66% StartBufferIO
+ 1.56% LockReleaseAll
+ 1.55% _bt_next
+ 0.78% LockAcquireExtended
+ 1.47% _bt_next
+ 0.75% _bt_relandgetbuf
to me. Now that's with the client count 496, but it's similar with lower
counts.
BTW, that profile *clearly* indicates we should make StrategyGetBuffer()
smarter.

Which is nearly trivial now that atomics are in. Check out the attached
WIP patch which eliminates the spinlock from StrategyGetBuffer() unless
there's buffers on the freelist.

Is this safe?
+ victim = pg_atomic_fetch_add_u32(&StrategyControl->nextVictimBuffer, 1);
- if (++StrategyControl->nextVictimBuffer >= NBuffers)
+ buf = &BufferDescriptors[victim % NBuffers];
+
+ if (victim % NBuffers == 0)
<snip>
I don't think there's any guarantee you could keep nextVictimBuffer
from wandering off the end.

Not sure what you mean? It'll cycle around at 2^32. The code doesn't try
to avoid nextVictimBuffer from going above NBuffers. To avoid overrunning
&BufferDescriptors I'm doing % NBuffers.

Yes, that'll have the disadvantage of the first buffers being slightly
more likely to be hit, but for that to be relevant you'd need
unrealistically large amounts of shared_buffers.

Post by Merlin Moncure
You could buy a little safety by CAS'ing
zero in instead, followed by atomic increment. However that's still
pretty dodgy IMO and requires two atomic ops which will underperform
the spin in some situations.
I like Robert's idea to keep the spinlock but preallocate a small
amount of buffers, say 8 or 16.

I think that's pretty much orthogonal. I *do* think it makes sense to
increment nextVictimBuffer in bigger steps. But the above doesn't
prohibit doing so - and it'll still be be much better without the
spinlock. Same number of atomics, but no potential of spinning and no
potential of being put to sleep while holding the spinlock.

This callsite has a comparatively large likelihood of being put to sleep
while holding a spinlock because it can run for a very long time doing
nothing but spinlocking.

Greetings,

Andres Freund

[Merlin Moncure]

Post by Andres Freund

Post by Merlin Moncure

Post by Andres Freund
Which is nearly trivial now that atomics are in. Check out the attached
WIP patch which eliminates the spinlock from StrategyGetBuffer() unless
there's buffers on the freelist.

Is this safe?
+ victim = pg_atomic_fetch_add_u32(&StrategyControl->nextVictimBuffer, 1);
- if (++StrategyControl->nextVictimBuffer >= NBuffers)
+ buf = &BufferDescriptors[victim % NBuffers];
+
+ if (victim % NBuffers == 0)
<snip>
I don't think there's any guarantee you could keep nextVictimBuffer
from wandering off the end.

Not sure what you mean? It'll cycle around at 2^32. The code doesn't try
to avoid nextVictimBuffer from going above NBuffers. To avoid overrunning
&BufferDescriptors I'm doing % NBuffers.
Yes, that'll have the disadvantage of the first buffers being slightly
more likely to be hit, but for that to be relevant you'd need
unrealistically large amounts of shared_buffers.

Right -- my mistake. That's clever. I think that should work well.

Post by Andres Freund
I think that's pretty much orthogonal. I *do* think it makes sense to
increment nextVictimBuffer in bigger steps. But the above doesn't
prohibit doing so - and it'll still be be much better without the
spinlock. Same number of atomics, but no potential of spinning and no
potential of being put to sleep while holding the spinlock.
This callsite has a comparatively large likelihood of being put to sleep
while holding a spinlock because it can run for a very long time doing
nothing but spinlocking.

A while back, I submitted a minor tweak to the clock sweep so that,
instead of spinlocking every single buffer header as it swept it just
did a single TAS as a kind of a trylock and punted to the next buffer
if the test failed on the principle there's not good reason to hang
around. You only spin if you passed the first test; that should
reduce the likelihood of actual spinning to approximately zero. I
still maintain there's no reason not to do that (I couldn't show a
benefit but that was because mapping list locking was masking any
clock sweep contention at that time).

merlin

[Amit Kapila]

Post by Merlin Moncure
A while back, I submitted a minor tweak to the clock sweep so that,
instead of spinlocking every single buffer header as it swept it just
did a single TAS as a kind of a trylock and punted to the next buffer
if the test failed on the principle there's not good reason to hang
around. You only spin if you passed the first test; that should
reduce the likelihood of actual spinning to approximately zero. I
still maintain there's no reason not to do that (I couldn't show a
benefit but that was because mapping list locking was masking any
clock sweep contention at that time).

If you feel that can now show the benefit, then I think you can rebase
it for the coming commit fest (which is going to start today).


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Andres Freund

Post by Andres Freund

Post by Andres Freund

Post by Amit Kapila
Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all the
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client

count,

Post by Andres Freund

Post by Amit Kapila
however after that it starts dipping (in above config unto ~4.4%).

Hm. Interesting. I don't see that dip on x86.

Is it possible that implementation of some atomic operation is costlier
for particular architecture?

Yes, sure. And IIRC POWER improved atomics performance considerably for
POWER8...

Post by Andres Freund
I have tried again for scale factor 3000 and could see the dip and this
time I have even tried with 175 client count and the dip is

approximately

Post by Andres Freund

Post by Andres Freund
5% which is slightly more than 160 client count.

For my tests on Power8, the profile looks somewhat similar to below
profile mentioned by you, please see this mail:
http://www.postgresql.org/message-id/***@mail.gmail.com

However on Power7, the profile looks different which I have
posted above thread.

Post by Andres Freund
BTW, that profile *clearly* indicates we should make StrategyGetBuffer()
smarter.

Yeah, even bgreclaimer patch is able to achieve the same, however
after that the contention moves to somewhere else as you can see
in above link.

Post by Andres Freund

Post by Andres Freund
Here it goes..
Lwlock_contention patches - client_count=128
----------------------------------------------------------------------
+ 7.95% postgres postgres [.] GetSnapshotData
+ 3.58% postgres postgres [.] AllocSetAlloc
+ 2.51% postgres postgres [.] _bt_compare
+ 2.44% postgres postgres [.]
hash_search_with_hash_value
+ 2.33% postgres [kernel.kallsyms] [k] .__copy_tofrom_user
+ 2.24% postgres postgres [.] AllocSetFreeIndex
+ 1.75% postgres postgres [.]
pg_atomic_fetch_add_u32_impl

Uh. Huh? Normally that'll be inline. That's compiled with gcc? What were
the compiler settings you used?

Nothing specific, for performance tests where I have to take profiles
I use below:
./configure --prefix=<installation_path> CFLAGS="-fno-omit-frame-pointer"
make


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Andres Freund

Post by Andres Freund

Post by Andres Freund

Post by Amit Kapila
Observations
----------------------
a. The patch performs really well (increase upto ~40%) incase all

the

Post by Andres Freund

Post by Andres Freund

Post by Andres Freund

Post by Amit Kapila
data fits in shared buffers (scale factor -100).
b. Incase data doesn't fit in shared buffers, but fits in RAM
(scale factor -3000), there is performance increase upto 16 client

count,

Post by Andres Freund

Post by Amit Kapila
however after that it starts dipping (in above config unto ~4.4%).

Hm. Interesting. I don't see that dip on x86.

Is it possible that implementation of some atomic operation is costlier
for particular architecture?

Yes, sure. And IIRC POWER improved atomics performance considerably for
POWER8...

Post by Andres Freund
I have tried again for scale factor 3000 and could see the dip and this
time I have even tried with 175 client count and the dip is

approximately

Post by Andres Freund

Post by Andres Freund
5% which is slightly more than 160 client count.

I've run some short tests on hydra:

scale 1000:

base:
4GB:
tps = 296273.004800 (including connections establishing)
tps = 296373.978100 (excluding connections establishing)

8GB:
tps = 338001.455970 (including connections establishing)
tps = 338177.439106 (excluding connections establishing)

base + freelist:
4GB:
tps = 297057.523528 (including connections establishing)
tps = 297156.987418 (excluding connections establishing)

8GB:
tps = 335123.867097 (including connections establishing)
tps = 335239.122472 (excluding connections establishing)

base + LW_SHARED:
4GB:
tps = 296262.164455 (including connections establishing)
tps = 296357.524819 (excluding connections establishing)
8GB:
tps = 336988.744742 (including connections establishing)
tps = 337097.836395 (excluding connections establishing)

base + LW_SHARED + freelist:
4GB:
tps = 296887.981743 (including connections establishing)
tps = 296980.231853 (excluding connections establishing)

8GB:
tps = 345049.062898 (including connections establishing)
tps = 345161.947055 (excluding connections establishing)

I've also run some preliminary tests using scale=3000 - and I couldn't
see a performance difference either.

Note that all these are noticeably faster than your results.

Post by Amit Kapila

Post by Andres Freund

Post by Andres Freund
Lwlock_contention patches - client_count=128
----------------------------------------------------------------------
+ 7.95% postgres postgres [.] GetSnapshotData
+ 3.58% postgres postgres [.] AllocSetAlloc
+ 2.51% postgres postgres [.] _bt_compare
+ 2.44% postgres postgres [.]
hash_search_with_hash_value
+ 2.33% postgres [kernel.kallsyms] [k] .__copy_tofrom_user
+ 2.24% postgres postgres [.] AllocSetFreeIndex
+ 1.75% postgres postgres [.]
pg_atomic_fetch_add_u32_impl

Uh. Huh? Normally that'll be inline. That's compiled with gcc? What were
the compiler settings you used?

Nothing specific, for performance tests where I have to take profiles
./configure --prefix=<installation_path> CFLAGS="-fno-omit-frame-pointer"
make

Hah. Doing so overwrites the CFLAGS configure normally sets. Check
# CFLAGS are selected so:
# If the user specifies something in the environment, that is used.
# else: If the template file set something, that is used.
# else: If coverage was enabled, don't set anything.
# else: If the compiler is GCC, then we use -O2.
# else: If the compiler is something else, then we use -O, unless debugging.

so, if you do like above, you're compiling without optimizations... So,
include at least -O2 as well.

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund
tps = 296273.004800 (including connections establishing)
tps = 296373.978100 (excluding connections establishing)
tps = 338001.455970 (including connections establishing)
tps = 338177.439106 (excluding connections establishing)
tps = 297057.523528 (including connections establishing)
tps = 297156.987418 (excluding connections establishing)
tps = 335123.867097 (including connections establishing)
tps = 335239.122472 (excluding connections establishing)
tps = 296262.164455 (including connections establishing)
tps = 296357.524819 (excluding connections establishing)
tps = 336988.744742 (including connections establishing)
tps = 337097.836395 (excluding connections establishing)
tps = 296887.981743 (including connections establishing)
tps = 296980.231853 (excluding connections establishing)
tps = 345049.062898 (including connections establishing)
tps = 345161.947055 (excluding connections establishing)
I've also run some preliminary tests using scale=3000 - and I couldn't
see a performance difference either.
Note that all these are noticeably faster than your results.

What is the client count?
Could you please post numbers you are getting for 3000 scale
factor for client count 128 and 175?

Post by Andres Freund

Post by Amit Kapila
Nothing specific, for performance tests where I have to take profiles
./configure --prefix=<installation_path>

CFLAGS="-fno-omit-frame-pointer"

Post by Andres Freund

Post by Amit Kapila
make

Hah. Doing so overwrites the CFLAGS configure normally sets. Check
# If the user specifies something in the environment, that is used.
# else: If the template file set something, that is used.
# else: If coverage was enabled, don't set anything.
# else: If the compiler is GCC, then we use -O2.
# else: If the compiler is something else, then we use -O, unless debugging.
so, if you do like above, you're compiling without optimizations... So,
include at least -O2 as well.

Hmm. okay, but is this required when we do actual performance
tests, because for that currently I don't use CFLAGS.


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Andres Freund
tps = 296273.004800 (including connections establishing)
tps = 296373.978100 (excluding connections establishing)
tps = 338001.455970 (including connections establishing)
tps = 338177.439106 (excluding connections establishing)
tps = 297057.523528 (including connections establishing)
tps = 297156.987418 (excluding connections establishing)
tps = 335123.867097 (including connections establishing)
tps = 335239.122472 (excluding connections establishing)
tps = 296262.164455 (including connections establishing)
tps = 296357.524819 (excluding connections establishing)
tps = 336988.744742 (including connections establishing)
tps = 337097.836395 (excluding connections establishing)
tps = 296887.981743 (including connections establishing)
tps = 296980.231853 (excluding connections establishing)
tps = 345049.062898 (including connections establishing)
tps = 345161.947055 (excluding connections establishing)
I've also run some preliminary tests using scale=3000 - and I couldn't
see a performance difference either.
Note that all these are noticeably faster than your results.

What is the client count?

160, because that was the one you reported the biggest regression.

Post by Amit Kapila
Could you please post numbers you are getting for 3000 scale
factor for client count 128 and 175?

Yes, although not tonight.... Also from hydra?

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
Nothing specific, for performance tests where I have to take profiles
./configure --prefix=<installation_path>

CFLAGS="-fno-omit-frame-pointer"

Post by Andres Freund

Post by Amit Kapila
make

Hah. Doing so overwrites the CFLAGS configure normally sets. Check
# If the user specifies something in the environment, that is used.
# else: If the template file set something, that is used.
# else: If coverage was enabled, don't set anything.
# else: If the compiler is GCC, then we use -O2.
# else: If the compiler is something else, then we use -O, unless

debugging.

Post by Andres Freund
so, if you do like above, you're compiling without optimizations... So,
include at least -O2 as well.

Hmm. okay, but is this required when we do actual performance
tests, because for that currently I don't use CFLAGS.

I'm not sure what you mean? You need to include -O2 in CFLAGS whenever
you override it. Your profile was clearly without inlining... And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
Could you please post numbers you are getting for 3000 scale
factor for client count 128 and 175?

Yes, although not tonight....

No issues, whenever you get it.

Post by Andres Freund
Also from hydra?

Yes. One more thing I would like to share with you is that while doing
tests, there are some other settings change in postgresql.conf

maintenance_work_mem = 1GB
synchronous_commit = off
wal_writer_delay = 20ms
checkpoint_segments=256
checkpoint_timeout =35min

I don't think these parameters matter for the tests we are doing, but
still I thought it is good to share, because I forgot to send some of
these non-default settings in previous mail.

Post by Andres Freund

Post by Amit Kapila

Post by Andres Freund

Post by Amit Kapila
Nothing specific, for performance tests where I have to take profiles
./configure --prefix=<installation_path>

CFLAGS="-fno-omit-frame-pointer"

Post by Andres Freund

Post by Amit Kapila
make

Hah. Doing so overwrites the CFLAGS configure normally sets. Check
# If the user specifies something in the environment, that is used.
# else: If the template file set something, that is used.
# else: If coverage was enabled, don't set anything.
# else: If the compiler is GCC, then we use -O2.
# else: If the compiler is something else, then we use -O, unless

debugging.

Post by Andres Freund
so, if you do like above, you're compiling without optimizations... So,
include at least -O2 as well.

Hmm. okay, but is this required when we do actual performance
tests, because for that currently I don't use CFLAGS.

I'm not sure what you mean? You need to include -O2 in CFLAGS whenever
you override it.

okay, thats what I wanted to ask you, so that we should not see different
numbers due to the way code is built.

When I do performance tests where I don't want to see profile,
I use below statement:
./configure --prefix=<installation_path>

Post by Andres Freund
And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

You have reported numbers at 1000 scale factor and mine were
at 3000 scale factor, so I think the difference is expected.

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Hi,

Post by Amit Kapila

Post by Andres Freund
And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

You have reported numbers at 1000 scale factor and mine were
at 3000 scale factor, so I think the difference is expected.

The numbers for 3000 show pretty much the same:

SCALE 128 160 175
HEAD 352113 339005 336491
LW_SHARED 365874 347931 342528

Hm. I wonder if you're using pgbench without -M prepared? That'd about
explain the difference.

Greetings,

Andres Freund

[Amit Kapila]

Post by Amit Kapila

Post by Andres Freund
And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

You have reported numbers at 1000 scale factor and mine were
at 3000 scale factor, so I think the difference is expected.

SCALE 128 160 175
HEAD 352113 339005 336491
LW_SHARED 365874 347931 342528
Hm. I wonder if you're using pgbench without -M prepared?

No, I use below statement:
./pgbench -c 128 -j 128 -T 300 -S -M prepared postgres

That'd about
explain the difference.

Here I think first thing to clarify is why the numbers on HEAD are
different? Another thing is that I generally see difference in
numbers at 1000 and 3000 scale factor (although I have not run
lately), but in your case the numbers are almost same.

I will try once more by cleaning every thing(installation, data_dir, etc..)
but not today...

With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Amit Kapila

Post by Amit Kapila

Post by Andres Freund
And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

You have reported numbers at 1000 scale factor and mine were
at 3000 scale factor, so I think the difference is expected.

SCALE 128 160 175
HEAD 352113 339005 336491
LW_SHARED 365874 347931 342528
Hm. I wonder if you're using pgbench without -M prepared?

./pgbench -c 128 -j 128 -T 300 -S -M prepared postgres

That'd about
explain the difference.

Here I think first thing to clarify is why the numbers on HEAD are
different?

I have taken the latest code and recreated the database and tried
again on power-7 m/c (hydra) and below is the result:

Result with -M prepared:
Duration of each individual run - 5 mins

HEAD – commit 494affb

Shared_buffers=8GB; Scale Factor = 3000

Client Count/No. Of Runs (tps) 128 160 Run-1 258385 239908 Run-2 257835
238624 Run-3 255967 237905

Result without -M prepared:
Duration of each individual run - 5 mins


HEAD – commit 494affb

Shared_buffers=8GB; Scale Factor = 3000

Client Count/No. Of Runs (tps) 128 160 Run-1 228747 220961 Run-2 229817
214464 Run-3 227386 216619

I am not sure why we are seeing difference even though running
on same m/c with same configuration. I think there is some
difference in the way we are running tests, if you don't mind
could you please share the exact steps and non-default postgresql.conf
settings with me. The below list of things could be useful for
me to reproduce the numbers you are seeing:

a. build steps (any script you are using)
b. non-default postgresql.conf settings
c. Exact pgbench statements used


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Amit Kapila]

Post by Amit Kapila
I am not sure why we are seeing difference even though running
on same m/c with same configuration.

I have tried many times, but I could not get the numbers you have
posted above with HEAD, however now trying with the latest version
[1] posted by you, everything seems to be fine at this workload.
The data at higher client count is as below:

HEAD – commit 494affb

Shared_buffers=8GB; Scale Factor = 3000

Client Count/No. Of Runs (tps) 64 128 Run-1 271799 247777 Run-2 274341
245207 Run-3 275019 252258





HEAD – commit 494affb + wait free lw_shared_v2

Shared_buffers=8GB; Scale Factor = 3000

Client Count/No. Of Runs (tps) 64 128 Run-1 286209 274922 Run-2 289101
274495 Run-3 289639 273633

So I am planning to proceed further with the review/test of your
latest patch.

According to me, below things are left from myside:
a. do some basic tpc-b tests with patch
b. re-review latest version posted by you

I know that you have posted optimization into StrategyGetBuffer() in
this thread, however I feel we can evaluate it separately unless you
are of opinion that both the patches should go together.

[1]
http://www.postgresql.org/message-id/***@alap3.anarazel.de



With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Amit Kapila

Post by Amit Kapila
I am not sure why we are seeing difference even though running
on same m/c with same configuration.

I have tried many times, but I could not get the numbers you have
posted above with HEAD, however now trying with the latest version
[1] posted by you, everything seems to be fine at this workload.

I'll try to reproduce it next week. But I don't think it matters all
that much. Personally so far the performance numbers don't seem to
indicate much reason to wait any further. We sure improve further, but I
don't see much reason to wait because of that.

Post by Amit Kapila
HEAD – commit 494affb
Shared_buffers=8GB; Scale Factor = 3000
Client Count/No. Of Runs (tps) 64 128 Run-1 271799 247777 Run-2 274341
245207 Run-3 275019 252258
HEAD – commit 494affb + wait free lw_shared_v2
Shared_buffers=8GB; Scale Factor = 3000
Client Count/No. Of Runs (tps) 64 128 Run-1 286209 274922 Run-2 289101
274495 Run-3 289639 273633

So here the results with LW_SHARED were consistently better, right? You
saw performance degradations here earlier?

Post by Amit Kapila
So I am planning to proceed further with the review/test of your
latest patch.
a. do some basic tpc-b tests with patch
b. re-review latest version posted by you

Cool!

Post by Amit Kapila
I know that you have posted optimization into StrategyGetBuffer() in
this thread, however I feel we can evaluate it separately unless you
are of opinion that both the patches should go together.
[1]

No, I don't think they should go together - I wrote that patch because
it was the bottleneck in the possibly regressing test and I wanted to
see the full effect. Although I do think we should apply it ;)

Greetings,

Andres Freund

[Amit Kapila]

Post by Andres Freund

Post by Amit Kapila
HEAD – commit 494affb + wait free lw_shared_v2
Shared_buffers=8GB; Scale Factor = 3000
Client Count/No. Of Runs (tps) 64 128 Run-1 286209 274922 Run-2 289101
274495 Run-3 289639 273633

So here the results with LW_SHARED were consistently better, right?

Yes.

Post by Andres Freund
You
saw performance degradations here earlier?

Yes.

Post by Andres Freund

Post by Amit Kapila
So I am planning to proceed further with the review/test of your
latest patch.
a. do some basic tpc-b tests with patch

I have done few tests, the results of which are below, the data indicates
that neither there is any noticeable gain nor any noticeable loss on tpc-b
tests which I think is what could have been expected of this patch.
There is slight variation at few client counts (for sync_commit =off,
at 32 and 128), however I feel that is just noise as I don't see any
general trend.

Performance Data
----------------------------
IBM POWER-8 24 cores, 192 hardware threads
RAM = 492GB
Database Locale =C
max_connections =300
checkpoint_segments=300
checkpoint_timeout =15min
maintenance_work_mem = 1GB
checkpoint_completion_target = 0.9
Client Count = number of concurrent sessions and threads (ex. -c 8 -j 8)
Duration of each individual run = 30mins
Test mode - tpc-b

Below data is median of 3 runs, detailed data is attached with this
mail.


Scale_factor =3000; shared_buffers=8GB;

Patch/Client_count 8 16 32 64 128 HEAD 3849 4889 3569 3845 4547
LW_SHARED 3844 4787 3532 3814 4408

Scale_factor =3000; shared_buffers=8GB; synchronous_commit=off;

Patch/Client_count 8 16 32 64 128 HEAD 5966 8297 10084 9348 8836
LW_SHARED 6070 8612 8839 9503 8584

While doing performance tests, I noticed a hang at higher client
counts with patch. I have tried to check call stack for few of
processes and it is as below:

#0 0x0000008010933e54 in .semop () from /lib64/libc.so.6
#1 0x0000000010286e48 in .PGSemaphoreLock ()
#2 0x00000000102f68bc in .LWLockAcquire ()
#3 0x00000000102d1ca0 in .ReadBuffer_common ()
#4 0x00000000102d2ae0 in .ReadBufferExtended ()
#5 0x00000000100a57d8 in ._bt_getbuf ()
#6 0x00000000100a6210 in ._bt_getroot ()
#7 0x00000000100aa910 in ._bt_search ()
#8 0x00000000100ab494 in ._bt_first ()
#9 0x00000000100a8e84 in .btgettuple ()
..

#0 0x0000008010933e54 in .semop () from /lib64/libc.so.6
#1 0x0000000010286e48 in .PGSemaphoreLock ()
#2 0x00000000102f68bc in .LWLockAcquire ()
#3 0x00000000102d1ca0 in .ReadBuffer_common ()
#4 0x00000000102d2ae0 in .ReadBufferExtended ()
#5 0x00000000100a57d8 in ._bt_getbuf ()
#6 0x00000000100a6210 in ._bt_getroot ()
#7 0x00000000100aa910 in ._bt_search ()
#8 0x00000000100ab494 in ._bt_first ()
...

The test configuration is as below:
Test env - Power - 7 (hydra)
scale_factor - 3000
shared_buffers - 8GB
test mode - pgbench read only

test execution -
./pgbench -c 128 -j 128 -T 1800 -S -M prepared postgres

I have ran it for half an hour, but it doesn't came out even after
~2 hours. It doesn't get reproduced every time, currently I am
able to reproduce it and the m/c is in same state, if you want any
info, let me know (unfortunately binaries are in release mode, so
might not get enough information).

Post by Andres Freund

Post by Amit Kapila
b. re-review latest version posted by you

Cool!

I will post my feedback for code separately, once I am able to
completely review the new versions.


With Regards,
Amit Kapila.
EnterpriseDB: http://www.enterprisedb.com

[Andres Freund]

Post by Andres Freund
Hi,

Post by Amit Kapila

Post by Andres Freund
And since
your general performance numbers are a fair bit lower than what I see
with, hopefully, the same code on the same machine...

You have reported numbers at 1000 scale factor and mine were
at 3000 scale factor, so I think the difference is expected.

SCALE 128 160 175
HEAD 352113 339005 336491
LW_SHARED 365874 347931 342528
Hm. I wonder if you're using pgbench without -M prepared? That'd about
explain the difference.

Started a test for a run without -M prepared (i.e. -M simple):

SCALE 1 2 4 8 16 32 64 96 128 160 196
HEAD 7968 15132 31147 63395 123551 180436 242098 263625 249652 244240 232679
LW_SHARED 8268 15032 31267 63911 118943 180447 247067 269262 259165 247166 231292

This really doesn't look exciting to me.

scale 200, -M prepared:
SCALE 1 2 4 8 16 32 64 96 128 160 196
HEAD 13032 24712 50967 103801 201745 279226 380844 392623 379953 357206 361072
LW_SHARED 12997 25134 51119 102920 199597 282511 422424 460222 447662 436959 418519

My guess is that the primary benefit on systems with few sockets, like
this, is that with the new code there's far fewer problems with
processes sleeping (being scheduled out) while holding a spinlock.

Greetings,

Andres Freund