You can listen to me talk to Carl & Richard on RavenDB Sharding here.

What is data sharding, and why do you need it? Carl and Richard talk to Oren Eini about his latest work on RavenDB, including the new data sharding feature. Oren talks about the power of sharding a database across multiple servers to improve performance on massive data sets. While a sharded database is typically in a single data center, it is possible to distribute the shards across multiple locations. The conversation explores the advantages and disadvantages of the different approaches, including that you might not need it today, but it's great to know it's there when you do!

This episode was recorded a while ago, and just went live.

In the previous post, I showed a very simple request router that would always select the fastest node. That worked for a long while, until it didn’t, and the challenge is figuring out why.

As it turns out, the issue is a simple one of spooky action at a distance. Here is what happens. Assume that we have three servers and 10 clients. Each server is sized to handle 4 clients. So far, so good, the system has the capacity to spare.

The problem is in the manner in which clients will detect which is the fastest node in the cluster. The only thing that is considered is the state of the node at the time of selection. At that time, we may end up with all the nodes selecting one particular node as the fastest.

In other words, we have three servers, two of them have no clients talking to them and one of the servers has all the clients talking to it. That results in that node going down, obviously. The clients would then react appropriately, and select a new node to talk to. All of them would do that, find the fastest node, and… bring it down as well. Rinse & repeat.

The issue can be stated as Time Of Check vs Time Of Use, but also as a race condition, where all individual nodes end up doing a synchronized “wave” operation that kills the system.

How do you prevent this?

You introduce randomness into the system. You don’t test the status once, but re-check on a regular basis so you can respond to shifting load. You should also introduce randomness into the process. So the nodes won’t all do this exactly at the same time and end up in the same position.

Side note: Current state in Israel right now is bad. I’m writing this blog post as a form of escapism so I can talk about something that makes sense and follow logic and reason. I’ll not comment on the current status otherwise in this area.

Consider the following scenario. We have a bunch of servers and clients. The clients want to send requests for processing to the fastest node that they have available. But the algorithm that was implemented has an issue, can you see what this is?

To simplify things, we are going to assume that the work that is being done for each request is the same, so we don’t need to worry about different request workloads.

The idea is that each client node will find the fastest node (usually meaning the nearest one) and if there is enough load on the server to have it start throwing errors, it will try to find another one. This system has successfully spread the load across all servers, until one day, the entire system went down. And then it stayed down.

Can you figure out what is the issue?

Today marks a very long journey for RavenDB as we release version 6.0 into the wild. This completes a multi-year effort on our part, which I’m really excited to share with you.

In version 6.0 the RavenDB team had a number of goals, with the primary one being, as always, providing a simple and powerful platform for your data and documents.

The full list of changes in the release is here and you can already download the binaries of the new release to try it out.

Below you’ll find more information about some of the highlights for the new release, but the executive summary for the technically inclined is:

• Deep integration with Kafka & RabbitMQ - allowing to send and consume messages directly from RavenDB.
• Sharding support allows you to run massively large databases easily and cheaply.
• Corax is a new indexing engine for RavenDB that can provide a 10x performance improvement for both indexing and queries.

With this new release we are also announcing a pricing update for our commercial customers. You can find more details at the bottom of this post.

Before we get to the gist of the changes, I want to spend a minute talking about this release. One of the reasons that we took so long to craft this release was an uncompromising eye toward ease of use, performance, and quality of experience.

It may sound strange to talk about “quality of experience” when discussing a database engine, we aren’t talking about a weekend trip or a romantic getaway, after all. The reason RavenDB exists is that we found other databases to be complicated and delicate beasts, and we wanted to create a database that would be a joy to work with. Something that you could throw your data and queries to and would just work.

With version 6.0, I’m happy to say that we managed to keep on doing just that while providing you with a host of new features and capabilities. Without further ado, allow me to tell you about the highlights of this release.

Sharding

Sharding is the process of splitting your data among multiple nodes, to allow you to scale beyond the terabyte range. RavenDB actually had this feature, as far back as 2010, when the 1.0 version came out. I’m speaking about this today because we have decided to revamp the entire approach to sharding in the new release. Sharding in RavenDB version 6.0 consists of selecting the sharded option on database creation time, and everything else is handled by the cluster. Splitting the data across the nodes, balancing load and giving you the feeling that you are working on a unified database is all our responsibility.

A chief goal of the new sharding feature in RavenDB was to make sure that you don’t have to be aware that you are running in a sharded environment. In fact, you can even switch the backend to a sharded database behind your application’s back, with no modifications to the application.

Your systems can take advantage of the sharding feature, but they aren’t forced to. This makes it easier when you need to scale because you don’t need an expensive rewrite or stop the world migration to move to the new architecture.

RavenDB, with or without sharding, offers you the same capabilities. Features such as indexing, map/reduce, ETL tasks, or subscriptions all work in the same way regardless of how you choose to store your data. A successful metric for us is that after you switch over to sharding, the only thing that you’ll notice is that you can now scale your systems even more. And, of course, you could do that easily, safely, and quickly.

I recorded a webinar showcasing our new sharding approach that you can watch.

Corax Indexing Engine

Corax indexing engine is now available in version 6.0. We have been working on that since 2014, and as you can imagine, it hasn’t been a simple -to-solve challenge. Corax has a single task, to be able to do everything that we use Lucene for, but do that faster. In fact, to do things much faster.

RavenDB has used Lucene as its indexing engine from the start, and Lucene has been the industry benchmark for search engines for a long time. In the context of RavenDB, we have optimized Lucene heavily for the past 15 years. When building Corax, it wasn’t sufficient to match what Lucene can do but to also significantly outperform it.

With version 6.0, we now offer both Corax and Lucene as indexing engines for RavenDB. You can choose to use either engine (globally or per index). Upgrading from an older version of RavenDB, you’ll keep on using the Lucene engine but can create new indexes with Corax.

Corax outperforms Lucene on a wide range of scenarios by an order of magnitude. This includes both indexing time and query latency. Corax also manages to be far faster than Lucene while utilizing less CPU and memory. One of the ways it does that is by preparing, in advance, the relevant data structures that Lucene needs to compute at runtime.

Corax indexes tend to consume about 10% - 20% more disk space than Lucene, as a trade for offering lowered memory usage, far faster querying performance, and reduced indexing time.

In the same manner as sharding, switching to the new indexing engine will get you performance improvements, with no difference in capabilities or features.

Kafka & RabbitMQ Integration

Kafka & RabbitMQ integration has been extended to version 6.0. RavenDB already supported writing to Kafka and RabbitMQ and we have now extended this capability to allow you to read messages from Kafka and RabbitMQ and turn them into documents.

This capability provides complete support for queuing infrastructure integration. RavenDB can work with existing pipelines and easily expose its data for the rest of the organization as well as consume messages from the queue and create or update documents accordingly.

Such documents are then available for querying, aggregation, etc. This makes it trivial to use a low-code solution to quickly and efficiently plug your system into Kafka and RabbitMQ. In the style of all RavenDB features, we take upon ourselves all the complexity inherent in such a solution. You only need to provide the connection string and the details about what you want to pull, and the full management of pulling or pushing data, handling distributed state, failover, and monitoring is the responsibility of the RavenDB cluster.

Data Archiving

Data archiving is a new capability in RavenDB, aiming to help users who deal with very large amounts of data. While a business wants to retain all its data forever, it typically works primarily with recent data. That is typically the data from the last year or two. As time goes by, old data accumulates, and going through irrelevant data consumes computing and memory resources from the database.

RavenDB version 6.0 offers a way to mark a document with an attribute, marking when RavenDB should move that document to archive mode. Aside from telling RavenDB at what time we should archive a document, there is nothing further that you need to do.

When the time comes to archive the document, RavenDB will change the document mode, which will have a number of effects. The document itself is retained, it is not deleted. If you aren’t already using document compression, the archived document is stored in a compressed manner, to reduce disk usage further.

Subscriptions and indexes will ignore archive documents and unless explicitly requested, archived documents will remain on the disk and consume no memory resources. Archiving reduces the resources consumed by archived documents while keeping them available if you need to look them up.

You can also construct indexing that would operate over archived and regular documents as well if you still want to allow queries over archived data. In this way, you can retain your entire dataset in the main system of record but still operate a significantly smaller chunk of that during normal operations.

Performance, observability, and usability enhancements have also been on the menu for the version 6.0 release. There are far too many individual features and improvements to count them all in this format. I encourage you to check the full release details on the website.

You are probably well aware of the drastic changes that happened in the business environment in the last few years. As part of adjusting to this new reality, we are updating our pricing to allow us to keep delivering top-notch solutions and support to our valued customers.

The new pricing policy is effective starting from January 1st, 2024.

What does this mean to you right now? Absolutely nothing until January 1st, 2024. All subscription renewals in 2023 keep their current price point.

An Early Birds benefit for our existing customers: renew your 2024 subscription while locking your 2023 price point.
* Available for Cloud and on-premises Customers.

In the weeks ahead, we'll provide detailed updates about these changes and ensure a smooth transition. Our goal is to empower you to make optimized choices about your RavenDB subscription. Your satisfaction remains our priority.

Our team is ready to assist you. Please reach out to us at sales@ravendb.net for any questions you have.

Finally, I would like to thank the RavenDB team for their hard work in delivering this version.

Our goal for this release was to deliver you a lot more while making sure that you’ll not be exposed to the underlying complexity of the problems that we are solving for you.

I believe that as you try out the new release, you’ll see how successful we are in providing you with an excellent database platform to take your systems to new heights.

And with that, I am left only with encouraging you to try out the new version. And let us know what you think about it.

The following bug cost me a bunch of time, can you see what I’m doing wrong?

For fun, it’s so nasty because usually, it will accidentally work.

In the previous post, I was able to utilize AVX to get some nice speedups. In general, I was able to save up to 57%(!) of the runtime in processing arrays of 1M items. That is really amazing, if you think about it. But my best effort only gave me a 4% improvement when using 32M items.

I decided to investigate what is going on in more depth, and I came up with the following benchmark. Given that I want to filter negative numbers, what would happen if the only negative number in the array was the first one?

In other words, let’s see what happens when we could write this algorithm as the following line of code:

array[1..].CopyTo(array);

The idea here is that we should measure the speed of raw memory copy and see how that compares to our code.

Before we dive into the results, I want to make a few things explicit. We are dealing here with arrays of long, when I’m talking about an array with 1M items, I’m actually talking about an 8MB buffer, and for the 32M items, we are talking about 256MB of memory.

I’m running these benchmarks on the following machine:

    AMD Ryzen 9 5950X 16-Core Processor

    Base speed:    3.40 GHz
     L1 cache:    1.0 MB
     L2 cache:    8.0 MB
     L3 cache:    64.0 MB

    Utilization    9%
     Speed    4.59 GHz

In other words, when we look at this, the 1M items (8MB) can fit into L2 (barely, but certainly backed by the L3 cache. For the 32M items (256MB), we are far beyond what can fit in the cache, so we are probably dealing with memory bandwidth issues.

I wrote the following functions to test it out:

Let’s look at what I’m actually testing here.

• CopyTo() – using the span native copying is the most ergonomic way to do things, I think.
• MemoryCopy() – uses a built-in unsafe API in the framework. That eventually boils down to a heavily optimized routine, which… calls to Memove() if the buffer overlaps (as they do in this case).
• MoveMemory() – uses a pinvoke to call to the Windows API to do the moving of memory for us.

Here are the results for the 1M case (8MB):

We can see some real surprises here. I’m using the FilterCmp (the very basic implementation) that I wrote.

I cannot explain why CopyTo() and MemoryMove() are so slow.

What is really impressive is that the FilterCmp_Avx() and MoveMemory() are so close in performance, and so much faster. To put it in another way, we are already at a stage where we are within shouting distance from the MoveMemory() performance. That is.. really impressive.

That said, what happens with 32M (256MB) ?

Now we are faster in the FilterCmp_Avx than MoveMemory. That is… a pretty big wow, and a really nice close for this blog post series, right? Except that we won’t be stopping here.

The way the task I set out works, we are actually filtering just the first item out, and then we are basically copying the memory. Let’s do some math: 256MB in 20.1ms means 12.4 GB/sec!

On this system, I have the following memory setup:

    64.0 GB

    Speed:    2133 MHz
     Slots used:    4 of 4
     Form factor:    DIMM
     Hardware reserved:    55.2 MB

I’m using DDR4 memory, so I can expect a maximum speed of 17GB/sec. In theory, I might be able to get more if the memory is located on different DIMMs, but for the size in question, that is not likely.

I’m going to skip the training montage of VTune, understanding memory architecture and figuring out what is actually going on.

Let’s drop everything and look at what we have with just AVX vs. MoveMemory:

The new baseline is MoveMemory, and in this run, we can see that the AVX code is slightly slower.

It’s sadly not uncommon to see numbers shift by those ranges when we are testing such micro-optimizations, mostly because we are subject to so many variables that can affect performance. In this case, I dropped all the other benchmarks, which may have changed things.

At any rate, using those numbers, we have 12.4GB/sec for MoveMemory() and 11.8GB/sec for the AVX version. The hardware maximum speed is 17GB/sec. So we are quite close to what can be done.

For that matter, I would like to point out that the trivial code completed the task in 11GB/sec, so that is quite respectable and shows that the issue here is literally getting the memory fast enough to the CPU.

Can we do something about that? I made a pretty small change to the AVX version, like so:

What are we actually doing here? Instead of loading the value and immediately using it, we are loading the next value, then we are executing the loop and when we iterate again, we will start loading the next value and process the current one. The idea is to parallelize load and compute at the instruction level.

Sadly, that didn’t seem to do the trick. I saw a 19% additional cost for that version compared to the vanilla AVX one on the 8MB run and a 2% additional cost on the 256MB run.

I then realized that there was one really important test that I had to also make, and wrote the following:

In other words, let's test the speed of moving memory and filling memory as fast as we possibly can. Here are the results:

This is really interesting, for a small buffer (8MB), MoveMemory is somehow faster. I don’t have a way to explain that, but it has been a pretty consistent result in my tests.

For the large buffer (256MB), we are seeing results that make more sense to me.

• MoveMemory – 12.5 GB / sec
• FIllMemory – 15.8 GB / sec

In other words, for MoveMemory, we are both reading and writing, so we are paying for memory bandwidth in both directions. For filling the memory, we are only writing, so we can get better performance (no need for reads).

In other words, we are talking about hitting the real physical limits of what the hardware can do. There are all sorts of tricks that one can pull, but when we are this close to the limit, they are almost always context-sensitive and dependent on many factors.

To conclude, here are my final results:

As you can see, just the AVX version is comparable or (slightly) beating the MoveMemory function.

I tried things like prefetching memory, both the next item, the next cache item and from the next page, using non-temporal load and stores and many other things, but this is a pretty tough challenge.

What is really interesting is that the original, very simple and obvious implementation, clocked at 11 GB/sec. After pulling pretty much all the stops and tricks, I was able to hit 12.5 GB / sec.

I don’t think anyone can look / update / understand the resulting code in any way without going through deep meditation. That is not a bad result at all. And along the way, I learned quite a bit about how the lowest level of the machine architecture is working.

In the previous post I discussed how we can optimize the filtering of negative numbers by unrolling the loop, looked into branchless code and in general was able to improve performance by up to 15% from the initial version we started with. We pushed as much as we could on what can be done using scalar code. Now it is the time to open a whole new world and see what we can do when we implement this challenge using vector instructions.

The key problem with such tasks is that SIMD, AVX and their friends were designed by… an interesting process using a perspective that makes sense if you can see in a couple of additional dimensions. I assume that at least some of that is implementation constraints, but the key issue is that when you start using SIMD, you realize that you don’t have general-purpose instructions. Instead, you have a lot of dedicated instructions that are doing one thing, hopefully well, and it is your role to compose them into something that would make sense. Oftentimes, you need to turn the solution on its head in order to successfully solve it using SIMD. The benefit, of course, is that you can get quite an amazing boost in speed when you do this.

The algorithm we use is basically to scan the list of entries and copy to the start of the list only those items that are positive. How can we do that using SIMD? The whole point here is that we want to be able to operate on multiple data, but this particular task isn’t trivial. I’m going to show the code first, then discuss what it does in detail:

We start with the usual check (if you’ll recall, that ensures that the JIT knows to elide some range checks, then we load the PremuteTable. For now, just assume that this is magic (and it is). The first interesting thing happens when we start iterating over the loop. Unlike before, now we do that in chunks of 4 int64 elements at a time. Inside the loop, we start by loading a vector of int64 and then we do the first odd thing. We call ExtractMostSignificantBits(), since the sign bit is used to mark whether a number if negative or not. That means that I can use a single instruction to get an integer with the bits set for all the negative numbers. That is particularly juicy for what we need, since there is no need for comparisons, etc.

If the mask we got is all zeroes, it means that all the numbers we loaded to the vector are positives, so we can write them as-is to the output and move to the next part. Things get interesting when that isn’t the case.

We load a permute value using some shenanigans (we’ll touch on that shortly) and call the PermuteVar8x32() method. The idea here is that we pack all the non-negative numbers to the start of the vector, then we write the vector to the output. The key here is that when we do that, we increment the output index only by the number of valid values.  The rest of this method just handles the remainder that does not fit into a vector.

The hard part in this implementation was to figure out how to handle the scenario where we loaded some negative numbers. We need a way to filter them, after all. But there is no SIMD instruction that allows us to do so. Luckily, we have the Avx2.PermuteVar8x32() method to help here. To confuse things, we don’t actually want to deal with 8x32 values. We want to deal with 4x64 values. There is Avx2.Permute4x64() method, and it will work quite nicely, with a single caveat. This method assumes that you are going to pass it a constant value. We don’t have such a constant, we need to be able to provide that based on whatever the masked bits will give us.

So how do we deal with this issue of filtering with SIMD? We need to move all the values we care about to the front of the vector. We have the method to do that, PermuteVar8x32() method, and we just need to figure out how to actually make use of this. PermuteVar8x32() accepts an input vector as well as a vector of the premutation you want to make. In this case, we are basing this on the 4 top bits of the 4 elements vector of int64. As such, there are a total of 16 options available to us. We have to deal with 32bits values rather than 64bits, but that isn’t that much of a problem.

Here is the premutation table that we’ll be using:

What you can see here is that when we have a 1 in the bits (shown in comments) we’ll not copy that to the vector. Let’s take a look at the entry of 0101, which may be caused by the following values [1,-2,3,-4].

When we look at the right entry at index #5 in the table: 2,3,6,7,0,0,0,0

What does this mean? It means that we want to put the 2nd int64 element in the source vector and move it as the first element of the destination vector, take the 3rd element from the source as the second element in the destination and discard the rest (marked as 0,0,0,0 in the table).

This is a bit hard to follow because we have to compose the value out of the individual 32 bits words, but it works quite well. Or, at least, it would work, but not as efficiently. This is because we would need to load the PermuteTableInts into a variable and access it, but there are better ways to deal with it. We can also ask the JIT to embed the value directly. The problem is that the pattern that the JIT recognizes is limited to ReadOnlySpan<byte>, which means that the already non-trivial int32 table got turned into this:

This is the exact same data as before, but using ReadOnlySpan<byte> means that the JIT can package that inside the data section and treat it as a constant value.

The code was heavily optimized, to the point where I noticed a JIT bug where these two versions of the code give different assembly output:

Here is what we get out:

This looks like an unintended consequence of Roslyn and the JIT each doing their (separate jobs), but not reaching the end goal. Constant folding looks like it is done mostly by Roslyn, but it does a scan like that from the left, so it wouldn’t convert $A * 4 * 8 to $A * 32. That is because it stopped evaluating the constants when it found a variable. When we add parenthesis, we isolate the value and now understand that we can fold it.

Speaking of assembly, here is the annotated assembly version of the code:

And after all of this work, where are we standing?

So it seems that the idea of using SIMD instruction has a lot of merit. Moving from the original code to the final version, we see that we can complete the same task in up to half the time.

I’m not quite sure why we aren’t seeing the same sort of performance on the 32M, but I suspect that this is likely because we far exceed the CPU cache and we have to fetch it all from memory, so that is as fast as it can go.

If you are interested in learning more, Lemire solves the same problem in SVE (SIMD for ARM) and Paul has a similar approach in Rust.

If you can think of further optimizations, I would love to hear your ideas.

In the previous post, we looked into what it would take to reduce the cost of filtering negative numbers. We got into the assembly and analyzed exactly what was going on. In terms of this directly, I don’t think that even hand-optimized assembly would take us further. Let’s see if there are other options that are available for us to get better speed.

The first thing that pops to mind here is to do a loop unrolling. After all, we have a very tight loop, if we can do more work per loop iteration, we might get better performance, no? Here is my first version:

And here are the benchmark results:

That is quite a jump, 6% – 11% savings is no joke. Let’s look at what is actually going on at the assembly level and see if we can optimize this further.

As expected, the code size is bigger, 264 bytes versus the 55 we previously got. But more importantly, we got the range check back, and a lot of them:

The JIT isn’t able to reason about our first for loop and see that all our accesses are within bounds, which leads to doing a lot of range checks, and likely slows us down. Even with that, we are still showing significant improvements here.

Let’s see what we can do with this:

With that, we expect to have no range checks and still be able to benefit from the unrolling.

That helped, by quite a lot, it seems, for most cases, the 32M items case, however, was slightly slower, which is quite a surprise.

Looking at the assembly, I can see that we still have branches, like so:

And here is why this is the case:

Now, can we do better here? It turns out that we can, by using a branchless version of the operation. Here is another way to write the same thing:

What happens here is that we are unconditionally setting the value in the array, but only increment if the value is greater than or equal to zero. That saves us in branches and will likely result in less code. In fact, let’s see what sort of assembly the JIT will output:

What about the performance? I decided to pit the two versions (normal and branchless) head to head and see what this will give us:

Surprisingly enough, it looks like the branchless version is very slightly slower. That is a surprise, since I would expect reducing the branches to be more efficient.

Looking at the assembly of those two, the branchless version is slightly bigger (10 bytes, not that meaningful). I think that the key here is that there is a 0.5% chance of actually hitting the branch, which is pretty low. That means that the branch predictor can likely do a really good job and we aren’t going to see any big benefits from the branchless version.

That said… what would happen if we tested that with 5% negatives? That difference in behavior may cause us to see a different result. I tried that, and the results were quite surprising. In the case of the 1K and 32M items, we see a slightl cost for the branchless version (additional 1% – 4%) while for the 1M entries there is an 18% reduction in latency for the branchless version.

I ran the tests again with a 15% change of negative, to see what would happen. In that case, we get:

As you can see, we have basically the same cost under 15% negatives for small values, a big improvement on the 1M scenario and not much improvement on the 32M scenario.

All in all, that is very interesting information. Digging into the exact why and how of that means pulling a CPU instruction profiler and starting to look at where we have stalls, which is a bit further that I want to invest in this scenario.

What if we’ll try to rearrange the code a little bit. The code looks like this (load the value and AddToOutput() immediately):

AddToOutput(ref itemsRef, Unsafe.Add(ref itemsRef, i + 0));

What if we split it a little bit, so the code will look like this:

The idea here is that we are trying to get the JIT / CPU to fetch the items before they are actually needed, so there would be more time for the memory to arrive.

Remember that for the 1M scenario, we are dealing with 8MB of memory and for the 32M scenario, we have 256MB. Here is what happens when we look at the loop prolog, we can see that it is indeed first fetching all the items from memory, then doing the work:

In terms of performance, that gives us a small win (1% – 2% range) for the 1M and 32M entries scenario.

The one last thing that I wanted to test is if we’ll unroll the loop even further, what would happen if we did 8 items per loop, instead of 4.

There is some improvement, (4% in the 1K scenario, 1% in the 32M scenario) but also slowdowns  (2% in the 1M scenario).

I think that this is probably roughly the end of the line as far as we can get for scalar code.

We already made quite a few strides in trying to parallelize the work the CPU is doing by just laying out the code as we would like it to be. We tried to control the manner in which it touches memory and in general, those are pretty advanced techniques.

To close this post, I would like to take a look at the gains we got. I’m comparing the first version of the code, the last version we had on the previous post and the unrolled version for both branchy and branchless with 8 operations at once and memory prefetching.

The unrolled 8 version is the winner by far, in this scenario (0.5% negatives). Since that is the scenario we have in the real code, that is what I’m focusing on.

Is there anything left to do here?

My next step is to explore whether using vector instructions will be a good option for us.

While working deep on the guts of RavenDB, I found myself with a seemingly simple task. Given a list of longs, I need to filter out all negative numbers as quickly as possible.

The actual scenario is that we run a speculative algorithm, given a potentially large list of items, we check if we can fulfill the request in an optimal fashion. However, if that isn’t possible, we need to switch to a slower code path that does more work.

Conceptually, this looks something like this:

That is the setup for this story. The problem we have now is that we now need to filter the results we pass to the RunManually() method.

There is a problem here, however. We marked the entries that we already used in the list by negating them. The issue is that RunManually() does not allow negative values, and its internal implementation is not friendly to ignoring those values.

In other words, given a Span<long>, I need to write the code that would filter out all the negative numbers. Everything else about the list of numbers should remain the same (the order of elements, etc).

From a coding perspective, this is as simple as:

Please note, just looking at this code makes me cringe a lot. This does the work, but it has an absolutely horrible performance profile. It allocates multiple arrays, uses a lambda, etc.

We don’t actually care about the entries here, so we are free to modify them without allocating a new value. As such, let’s see what kind of code we can write to do this work in an efficient manner. Here is what I came up with:

The way this works is that we scan through the list, skipping writing the negative lists, so we effectively “move down” all the non-negative lists on top of the negative ones. This has a cost of O(N) and will modify the entire array, the final output is the number of valid items that we have there.

In order to test the performance, I wrote the following harness:

We compare 1K, 1M and 32M elements arrays, each of which has about 0.5% negative, randomly spread across the range. Because we modify the values directly, we need to sprinkle the negatives across the array on each call. In this case, I’m testing two options for this task, one that uses a direct comparison (shown above) and one that uses bitwise or, like so:

I’m testing the cost of sprinkling negatives as well, since that has to be done before each benchmark call (since we modify the array during the call, we need to “reset” its state for the next one).

Given the two options, before we discuss the results, what would you expect to be the faster option? How would the size of the array matter here?

I really like this example, because it is simple, there isn’t any real complexity in what we are trying to do. And there is a very straightforward implementation that we can use as our baseline. That also means that I get to analyze what is going on at a very deep level. You might have noticed the disassembler attribute on the benchmark code, we are going to dive deep into that. For the same reason, we aren’t using exactly 1K, 1M, or 32M arrays, but slightly higher than that, so we’ll have to deal with remainders later on.

Let’s first look at what the JIT actually did here. Here is the annotated assembly for the FilterCmp function:

For the FilterOr, the code is pretty much the same, except that the key part is:

As you can see, the cmp option is slightly smaller, in terms of code size. In terms of performance, we have:

The costs are very close to one another, with Or being very slightly faster on low numbers, and Cmp being slightly faster on the larger sizes. Note that the difference level between them is basically noise. They have the same performance.

The question is, can we do better here?

Looking at the assembly, there is an extra range check in the main loop that the JIT couldn’t elide (the call to items[output++]). Can we do something about it, and would it make any difference in performance? Here is how I can remove the range check:

Here I’m telling the JIT: “I know what I’m doing”, and it shows.

Let’s look at the assembly changes between those two methods, first the prolog:

Here you can see what we are actually doing here. Note the last 4 instructions, we have a range check for the items, and then we have another check for the loop. The first will get you an exception, the second will just skip the loop. In both cases, we test the exact same thing. The JIT had a chance to actually optimize that, but didn’t.

Here is a funny scenario where adding code may reduce the amount of code generated. Let’s do another version of this method:

In this case, I added a check to handle the scenario of items being empty. What can the JIT do with this now? It turns out, quite a lot. We dropped 10 bytes from the method, which is a nice result of our diet.  Here is the annotated version of the assembly:

A lot of the space savings in this case come from just not having to do a range check, but you’ll note that we still do an extra check there (lines 12..13), even though we already checked that. I think that the JIT knows that the value is not zero at this point, but has to consider that the value may be negative.

If we’ll change the initial guard clause to: items.Length <= 0, what do you think will happen? At this point, the JIT is smart enough to just elide everything, we are at 55 bytes of code and it is a super clean assembly (not a sentence I ever thought I would use). I’ll spare you going through more assembly listing, but you can find the output here.

And after all of that, where are we at?

There is a very slight benefit to the NoRangeCheck, but even when we talk about 32M items, we aren’t talking about a lot of time.

The question what can we do better here?

Deep inside of the Corax indexing engine inside of RavenDB there is the notion of a posting list. A posting list is just an ordered set of entry ids that contains a particular term. During the indexing process, we need to add and remove items from that posting list. This ends up being something like this:

For fun, go and ask ChatGPT to write you the code for this task.

You can assume that there are no duplicates between the removals and additions, and that adding an existing item is a no-op (so just one value would be in the end result). Here is a quick solution for this task (not actually tested that much, mind, but sufficient to understand what I’m trying to do):

If you look at this code in terms of performance, you’ll realize that this is quite expensive. In terms of complexity, this is actually pretty good, we iterate over the arrays just once, and the number of comparisons is also bounded to the lengths of the list.

However, there is a big issue here, the number of branches that you have to deal with. Basically, every if and every for loop is going to add a tiny bit of latency to the system. This is because these are unpredictable branches, which are pretty nasty to deal with.

It turns out that the values that we put in the posting list are actually always a multiple of 4, so the bottom 2 bits are always cleared. That means that we actually have a different way to deal with it. Here is the new logic:

This code was written with an eye to being able to explain the algorithm, mind, not performance.

The idea goes like this. We flag the removals with a bit, then concatenate all the arrays together, sort them, and then do a single scan over the whole thing, removing duplicates and removals.

In the real code, we are using raw pointers, not a List, so there are no access checks, etc.

From an algorithmic perspective, this code makes absolutely no sense at all. We concatenate all the values together, then sort them (O(NlogN) operation) then scan it again?!

How can that be faster than a single scan across all three arrays? The answer is simple, we have a really efficient sort primitive (vxsort) that is able to sort things really fast (GB/sec). There is a really good series of posts that explain how that is achieved.

Since we consider sorting to be cheap, the rest of the work is just a single scan on the list, and there are no branches at all there. The code plays with the offset that we write into, figuring out whether we need to overwrite the current value (duplicate) or go back (removal), but in general it means that it can execute very quickly.

This approach also has another really important aspect. Take a look at the actual code that we have in production. This is from about an hour worth of profiling a busy indexing session:

And the more common code path:

In both of them, you’ll notice something really important. There isn’t a call to sorting at all in here. In fact, when I search for the relevant function, I find:

That is 25 ms out of over an hour.

How can this be? As efficient as the sorting can be, we are supposed to be calling it a lot.

Well, consider one scenario, what happens if:

• There are no removals
• All additions happen after the last existing item in the list

In this case, I don’t need to do anything beyond concatenate the lists. I can skip the entire process entirely, just copy the existing and additions to the output and call it a day.

Even when I do have a lot of removals and complicated merge processes, the code structure means that the CPU can get through this code very quickly. This isn’t super friendly for humans to read, but for the CPU, this is chump change.