To me transformation efficiency and distance losses actually seem to make 5-6% pretty impressive?
6,287 karma · joined April 14, 2015
andres [at] hn dot anarazel dot de
- Is the change architecturally right?
Particularly the latter LLMs seem still pretty useless at.
I invite you to start a discussion on the lists about that feature, I've wished for it before.
That's true only for the latter (and even then only at a isolation level that's not too strict).
I agree that there should be non-fork primitives, I'm just not that sure that performance is the best argument.
It's actually pretty easy to get compilers to use those, you mainly need a bunch of narrow accesses to neighboring memory. The oodle post contains a godbolt link to pretty ordinary c code triggering this.
I'd guess that you also need some other conditions (multiple in flight stores, high boost speeds) to trigger this.
It's not the common mode of deployment, but it's definitely in prod use.
> You can do it, but then you have to be able to retry all of your transactions, including read.
Pure read transactions shouldn't need to be retried in postgres due to serialization errors. You need to have read-write dependencies for that.
That's not to say that effectively read only transactions aren't affected by serializable, you do need to record the necessary metadata for the serialization logic to work.
FWIW, if you know your transaction is read only and long running, you can start a transaction with START TRANSACTION READ ONLY DEFERRABLE, which makes the start transaction slower, but then does not need to do any work related to serializable while the transaction is running.
It is able to? Configure huge_page_size=1GB?
Support for 2MB pages was added in 2014, for larger pages 2020.
Edit: year details.
My sleep has gotten so much better. I really didn't realize that alcohol didn't affect just the night after I had a drink, but even the next one or two nights..
Defaulting to throw-away-VMs for everything is also the right choice for something where the threat model includes attackers submitting patches/PRs. I'll never understand why folks were ok with just container separation for that (and often have no separation in runners).
It gets a bit worse with preempt_lazy - for me just 15% percent or so - because the lock holder is scheduled out a bit more often. But it was bad before.
> My question therefore was how come this regression hasn't been visible with huge pages turned off with older kernel versions? You say that it was but I can't find this data point.
I mean it wasn't a regression before, because this is how it has behaved for a long time.
This workload is not a realistic thing that anybody would encounter in this form in the real world. Even without the contention - which only happens the first time the buffer pool is filled - you lose so much by not using huge pages with a 100gb buffer pool that you will have many other issues.
We (postgres and me personally) were concerned enough about potential contention in this path that we did get rid of that lock half a year ago (buffer replacement selection has been lock free for close to a decade, just unused buffers were found via a list protected by this lock).
But the performance gains we saw were relatively small, we didn't measure large buffer pools without huge pages though.
And at least I didn't test with this many connections doing small random reads into a cold buffer pool, just because it doesn't seem that interesting.
Turns out to be pretty crucial for performance though... Not manipulating them with a single atomic leads to way way worse performance.
For quite a while it was a 32bit atomic, but I recently made it a 64bit one, to allow the content lock (i.e. protecting the buffer contents, rather than the buffer header) to be in the same atomic var. That's for one nice for performance, it's e.g. very common to release a pin and a lock at the same time and there are more fun perf things we can do in the future. But the real motivation was work on adding support for async writes - an exclusive locker might need to consume an IO completion for a write that's in flight that is prevent it from acquiring the lock. And that was hard to do with a separate content lock and buffer state...
> And there are like ten open coded spin waits around the uses... you certainly have my empathy :)
Well, nearly all of those are all to avoid needing to hold a spinlock, which, as lamented a lot around this issue, don't perform that well when really contended :)
We're on our way to barely ever need the spinlock for the buffer header, which then should allow us to get rid of many of those loops.
> This got me thinking about 64-bit futexes again. Obviously that can't work with PI... but for just FUTEX_WAIT/FUTEX_WAKE, why not?
It'd be pretty nice to have. There are lot of cases where one needs more lock state than one can really encode into a 32bit lock state.
I'm quite keen to experiment with the rseq time slice extension stuff. Think it'll help with some important locks (which are not spinlocks...).
> Now you've gotten me wondering. This issue is, in some sense, artificial: the actual conceptual futex unlock operation does not require sequential consistency. What's needed is (roughly, anyway) an release operation that synchronizes with whoever subsequently acquires the lock (on x86, any non-WC store is sufficient) along with a promise that the kernel will get notified eventually (and preferably fairly quickly) if there was a non-spinning sleeper. But there is no requirement that the notification occur in any particular order wrt anything else except that the unlock must be visible by the time the notification occurs [0]; there isn't even a requirement that the notification not occur if there is no futex waiter.
Hah.
> ... > But maybe there are sneaky tricks. I'm wondering whether CMPXCHG (no lock) is secretly good enough for this. Imagine a lock word where bit 0 set means locked and bit 1 set means that there is a waiter. The wait operation observes (via plain MOV?) that bit 0 is set and then sets bit 1 (let's say this is done with LOCK CMPXCHG for simplicity) and then calls futex_wait(), so it thinks the lock word has the value 3. The unlock operation does plain CMPXCHG to release the lock. The failure case would be that it reports success while changing the value from 1 to 0. I don't know whether this can happen on Intel or AMD architectures.
I suspect the problem isn't so much the lock prefix, but that the non-futex spinlock release just is a store, whereas a futex release has to be a RMW operation.
I'm talking out of my ass here, but my guess is that the reason for the performance gain of the plain-store-is-a-spinlock-release on x86 comes from being able to do the release via the store buffer, without having to wait for exclusive ownership of the cache line. Due to being a somewhat contended simple spinlock, often embedded on the same line as the to-be-protected data, it's common for the line not not be in modified ownership anymore at release.
A quick hack shows the contended performance to be nearly indistinguishable with a futex based lock. Which makes sense, non-PI futexes don't transfer the scheduler slice the lock owner, because they don't know who the lock owner is. Postgres' spinlock use randomized exponential backoff, so they don't prevent the lock owner from getting scheduled.
Thus the contention is worse with PREEMPT_LAZY, even with non-PI futexes (which is what typical lock implementations are based on), because the lock holder gets scheduled out more often.
Probably worth repeating: This contention is due to an absurd configuration that should never be used in practice.
On x86 a spinlock release doesn't need a memory barrier (unless you do insane things) / lock prefix, but a futex based lock does (because you otherwise may not realize you need to futex wake). Turns out that that increase in memory barriers causes regressions that are nontrivial to avoid.
Another difficulty is that most of the remaining spinlocks are just a single bit in a 8 larger byte atomic. Futexes still don't support anything but 4 bytes (we could probably get away with using it on a part of the 8 byte atomic with some reordering) and unfortunately postgres still supports platforms with no 8 byte atomics (which I think is supremely silly), and the support for a fallback implementation makes it harder to use futexes.
The spinlock triggering the contention in the report was just stupid and we only recently got around to removing it, because it isn't used during normal operation.
Edit: forgot to add that the spinlock contention is not measurable on much more extreme workloads when using huge pages. A 100GB buffer pool with 4KB pages doesn't make much sense.
With what we know so far, I expect that there are just about no real world workloads that aren't already completely falling over that will be affected.
> One possible reason is that the planner configuration can be different per connection, so the plans might not transfer
That's part of it, another big part is that the transactional DDL makes it more complicated, as different sessions might require different plans.