Benchmarking Commit Logs

In this article, we look at Apache Kafka and NATS Streaming, two messaging systems based on the idea of a commit log. We’ll compare some of the features of both but spend less time talking about Kafka since by now it’s quite well known. Similar to previous studies, we’ll attempt to quantify their general performance characteristics through careful benchmarking.

The purpose of this benchmark is to test drive the newly released NATS Streaming system, which was made generally available just in the last few months. NATS Streaming doesn’t yet support clustering, so we try to put its performance into context by looking at a similar configuration of Kafka.

Unlike conventional message queues, commit logs are an append-only data structure. This results in several nice properties like total ordering of messages, at-least-once delivery, and message-replay semantics. Jay Kreps’ blog post The Log is a great introduction to the concept and particularly why it’s so useful in the context of distributed systems and stream processing (his book I Heart Logs is an extended version of the blog post and is a quick read).

Kafka, which originated at LinkedIn, is by far the most popular and most mature implementation of the commit log (AWS offers their own flavor of it called Kinesis, and imitation is the sincerest form of flattery). It’s billed as a “distributed streaming platform for building real-time data pipelines and streaming apps.” The much newer NATS Streaming is actually a data-streaming layer built on top of Apcera’s high-performance publish-subscribe system NATS. It’s billed as “real-time streaming for Big Data, IoT, Mobile, and Cloud Native Applications.” Both have some similarities as well as some key differences.

Fundamental to the notion of a log is a way to globally order events. Neither NATS Streaming nor Kafka are actually a single log but many logs, each totally ordered using a sequence number or offset, respectively.

In Kafka, topics are partitioned into multiple logs which are then replicated across a number of servers for fault tolerance, making it a distributed commit log. Each partition has a server that acts as the leader. Cluster membership and leader election is managed by ZooKeeper.

NATS Streaming’s topics are called “channels” which are globally ordered. Unlike Kafka, NATS Streaming does not support replication or partitioning of channels, though my understanding is clustering support is slated for Q1 2017. Its message store is pluggable, so it can provide durability using a file-backed implementation, like Kafka, or simply an in-memory store.

NATS Streaming is closer to a hybrid of traditional message queues and the commit log. Like Kafka, it allows replaying the log from a specific offset, the beginning of time, or the newest offset, but it also exposes an API for reading from the log at a specific physical time offset, e.g. all messages from the last 30 seconds. Kafka, on the other hand, only has a notion of logical offsets (correction: Kafka added support for offset lookup by timestamp in . Generally, relying on physical time is an anti-pattern in distributed systems due to clock drift and the fact that clocks are not always monotonic. For example, imagine a situation where a NATS Streaming server is restarted and the clock is changed. Messages are still ordered by their sequence numbers but their timestamps might not reflect that. Developers would need to be aware of this while implementing their business logic.

With Kafka, it’s strictly on consumers to track their offset into the log (or the high-level consumer which stores offsets in ZooKeeper (correction: Kafka itself can now store offsets which is used by the new Consumer API, meaning clients do not have to manage offsets directly or rely on ZooKeeper)). NATS Streaming allows clients to either track their sequence number or use a durable subscription, which causes the server to track the last acknowledged message for a client. If the client restarts, the server will resume delivery starting at the earliest unacknowledged message. This is closer to what you would expect from a traditional message-oriented middleware like RabbitMQ.

Lastly, NATS Streaming supports publisher and subscriber rate limiting. This works by configuring the maximum number of in-flight (unacknowledged) messages either from the publisher to the server or from the server to the subscriber. Starting in version 0.9, Kafka supports a similar rate limiting feature that allows producer and consumer byte-rate thresholds to be defined for groups of clients with its Quotas protocol.

Kafka was designed to avoid tracking any client state on the server for performance and scalability reasons. Throughput and storage capacity scale linearly with the number of nodes. NATS Streaming provides some additional features over Kafka at the cost of some added state on the server. Since clustering isn’t supported, there isn’t really any scale or HA story yet, so it’s unclear how that will play out. That said, once replication is supported, there’s a lot of work going into verifying its correctness (which is a major advantage Kafka has).


Since NATS Streaming does not support replication at this time (0.3.1), we’ll compare running a single instance of it with file-backed persistence to running a single instance of Kafka ( We’ll look at both latency and throughput running on commodity hardware (m4.xlarge EC2 instances) with load generation and consumption each running on separate instances. In all of these benchmarks, the systems under test have not been tuned at all and are essentially in their “off-the-shelf” configurations.

We’ll first look at latency by publishing messages of various sizes, ranging from 256 bytes to 1MB, at a fixed rate of 50 messages/second for 30 seconds. Message contents are randomized to account for compression. We then plot the latency distribution by percentile on a logarithmic scale from the 0th percentile to the 99.9999th percentile. Benchmarks are run several times in an attempt to produce a “normalized” result. The benchmark code used is open source.

First, to establish a baseline and later get a feel for the overhead added by the file system, we’ll benchmark NATS Streaming with in-memory storage, meaning messages are not written to disk.

Unsurprisingly, the 1MB configuration has much higher latencies than the other configurations, but everything falls within single-digit-millisecond latencies.nats_mem

NATS Streaming 0.3.1 (in-memory persistence)

 Size 99% 99.9% 99.99% 99.999% 99.9999% 
256B 0.3750ms 1.0367ms 1.1257ms 1.1257ms 1.1257ms
1KB 0.38064ms 0.8321ms 1.3260ms 1.3260ms 1.3260ms
5KB 0.4408ms 1.7569ms 2.1465ms 2.1465ms 2.1465ms
1MB 6.6337ms 8.8097ms 9.5263ms 9.5263ms 9.5263ms

Next, we look at NATS Streaming with file-backed persistence. This provides the same durability guarantees as Kafka running with a replication factor of 1. By default, Kafka stores logs under /tmp. Many Unix distributions mount /tmp to tmpfs which appears as a mounted file system but is actually stored in volatile memory. To account for this and provide as level a playing field as possible, we configure NATS Streaming to also store its logs in /tmp.

As expected, latencies increase by about an order of magnitude once we start going to disk.


NATS Streaming 0.3.1 (file-backed persistence)

 Size 99% 99.9% 99.99% 99.999% 99.9999% 
256B 21.7051ms 25.0369ms 27.0524ms 27.0524ms 27.0524ms
1KB 20.6090ms 23.8858ms 24.7124ms 24.7124ms 24.7124ms
5KB 22.1692ms 35.7394ms 40.5612ms 40.5612ms 40.5612ms
1MB 45.2490ms 130.3972ms 141.1564ms 141.1564ms 141.1564ms

Since we will be looking at Kafka, there is an important thing to consider relating to fsync behavior. As of version 0.8, Kafka does not call fsync directly and instead relies entirely on the background flush performed by the OS. This is clearly indicated by their documentation:

We recommend using the default flush settings which disable application fsync entirely. This means relying on the background flush done by the OS and Kafka’s own background flush. This provides the best of all worlds for most uses: no knobs to tune, great throughput and latency, and full recovery guarantees. We generally feel that the guarantees provided by replication are stronger than sync to local disk, however the paranoid still may prefer having both and application level fsync policies are still supported.

However, NATS Streaming calls fsync every time a batch is written to disk by default. This can be disabled through the use of the –file_sync flag. By setting this flag to false, we put NATS Streaming’s persistence behavior closer in line with Kafka’s (again assuming a replication factor of 1).

As an aside, the comparison between NATS Streaming and Kafka still isn’t completely “fair”. Jay Kreps points out that Kafka relies on replication as the primary means of durability.

Kafka leaves [fsync] off by default because it relies on replication not fsync for durability, which is generally faster. If you don’t have replication I think you probably need fsync and maybe some kind of high integrity file system.

I don’t think we can provide a truly fair comparison until NATS Streaming supports replication, at which point we will revisit this.

To no one’s surprise, setting –file_sync=false has a significant impact on latency, shown in the distribution below.


In fact, it’s now in line with the in-memory performance as before for 256B, 1KB, and 5KB messages, shown in the comparison below.


For a reason I have yet to figure out, the latency for 1MB messages is roughly an order of magnitude faster when fsync is enabled after the 95th percentile, which seems counterintuitive. If anyone has an explanation, I would love to hear it. I’m sure there’s a good debug story there. The distribution below shows the 1MB configuration for NATS Streaming with and without fsync enabled and just how big the difference is at the 95th percentile and beyond.


NATS Streaming 0.3.1 (file-backed persistence, –file_sync=false)

 Size 99% 99.9% 99.99% 99.999% 99.9999% 
256B 0.4304ms 0.8577ms 1.0706ms 1.0706ms 1.0706ms
1KB 0.4372ms 1.5987ms 1.8651ms 1.8651ms 1.8651ms
5KB 0.4939ms 2.0828ms 2.2540ms 2.2540ms 2.2540ms
1MB 1296.1464ms 1556.1441ms 1596.1457ms 1596.1457ms 1596.1457ms

Kafka with replication factor 1 tends to have higher latencies than NATS Streaming with –file_sync=false. There was one potential caveat here Ivan Kozlovic pointed out to me in that NATS Streaming uses a caching optimization for reads that may put it at an advantage.

Now, there is one side where NATS Streaming *may* be looking better and not fair to Kafka. By default, the file store keeps everything in memory once stored. This means look-ups will be fast. There is only a all-or-nothing mode right now, which means either cache everything or nothing. With caching disabled (–file_cache=false), every lookup will result in disk access (which when you have 1 to many subscribers will be bad). I am working on changing that. But if you do notice that in Kafka, consuming results in a disk read (given the other default behavior described above, they actually may not ;-)., then you could disable NATS Streaming file caching.

Fortunately, we can verify if Kafka is actually going to disk to read messages back from the log during the benchmark using iostat. We see something like this for the majority of the benchmark duration:

avg-cpu:  %user   %nice %system %iowait  %steal   %idle
          13.53    0.00   11.28    0.00    0.00   75.19

Device:    tps   Blk_read/s   Blk_wrtn/s   Blk_read   Blk_wrtn
xvda      0.00         0.00         0.00          0          0

Specifically, we’re interested in Blk_read, which indicates the total number of blocks read. It appears that Kafka does indeed make heavy use of the operating system’s page cache as Blk_wrtn and Blk_read rarely show any activity throughout the entire benchmark. As such, it seems fair to leave NATS Streaming’s –file_cache=true, which is the default.

One interesting point is Kafka offloads much of its caching to the page cache and outside of the JVM heap, clearly in an effort to minimize GC pauses. I’m not clear if the cache Ivan refers to in NATS Streaming is off-heap or not (NATS Streaming is written in Go which, like Java, is a garbage-collected language).

Below is the distribution of latencies for 256B, 1KB, and 5KB configurations in Kafka.


Similar to NATS Streaming, 1MB message latencies tend to be orders of magnitude worse after about the 80th percentile. The distribution below compares the 1MB configuration for NATS Streaming and Kafka.


Kafka (replication factor 1)

 Size 99% 99.9% 99.99% 99.999% 99.9999% 
256B 0.9230ms 1.4575ms 1.6596ms 1.6596ms 1.6596ms
1KB 0.5942ms 1.3123ms 17.6556ms 17.6556ms 17.6556ms
5KB 0.7203ms 5.7236ms 18.9334ms 18.9334ms 18.9334ms
1MB 5337.3174ms 5597.3315ms 5617.3199ms 5617.3199ms 5617.3199ms

The percentile distributions below compare NATS Streaming and Kafka for the 256B, 1KB, and 5KB configurations, respectively.




Next, we’ll look at overall throughput for the two systems. This is done by publishing 100,000 messages using the same range of sizes as before and measuring the elapsed time. Specifically, we measure throughput at the publisher and the subscriber.

Despite using an asynchronous publisher in both the NATS Streaming and Kafka benchmarks, we do not consider the publisher “complete” until it has received acks for all published messages from the server. In Kafka, we do this by setting request.required.acks to 1, which means the leader replica has received the data, and consuming the received acks. This is important because the default value is 0, which means the producer never waits for an ack from the broker. In NATS Streaming, we provide an ack callback on every publish. We use the same benchmark configuration as the latency tests, separating load generation and consumption on different EC2 instances. Note the log scale in the following charts.

Once again, we’ll start by looking at NATS Streaming using in-memory persistence. The truncated 1MB send and receive throughputs are 93.01 messages/second.


For comparison, we now look at NATS Streaming with file persistence and –file_sync=false. As before, this provides the closest behavior to Kafka’s default flush behavior. The second chart shows a side-by-side comparison between NATS Streaming with in-memory and file persistence.



Lastly, we look at Kafka with replication factor 1. Throughput significantly deteriorates when we set request.required.acks = 1 since the producer must wait for all acks from the server. This is important though because, by default, the client does not require an ack from the server. If this were the case, the producer would have no idea how much data actually reached the server once it finished—it could simply be buffered in the client, in flight over the wire, or in the server but not yet on disk. Running the benchmark with request.required.acks = 0 yields much higher throughput on the sender but is basically an exercise in how fast you can write to a channel using the Sarama Go client—slightly misleading.


Looking at some comparisons of Kafka and NATS Streaming, we can see that NATS Streaming has higher throughput in all but a few cases.



I want to repeat the disclaimer from before: the purpose of this benchmark is to test drive the newly released NATS Streaming system (which as mentioned earlier, doesn’t yet support clustering), and put its performance into context by looking at a similar configuration of Kafka.

Kafka generally scales very well, so measuring the throughput of a single broker with a single producer and single consumer isn’t particularly meaningful. In reality, we’d be running a cluster with several brokers and partitioning our topics across them.

For as young as it is, NATS Streaming has solid performance (which shouldn’t come as much of a surprise considering the history of NATS itself), and I imagine it will only get better with time as the NATS team continues to optimize. In some ways, NATS Streaming bridges the gap between the commit log as made popular by Kafka and the conventional message queue as made popular by protocols like JMS, AMQP, STOMP, and the like.

The bigger question at this point is how NATS Streaming will tackle scaling and replication (a requirement for true production-readiness in my opinion). Kafka was designed from the ground up for high scalability and availability through the use of external coordination (read ZooKeeper). Naturally, there is a lot of complexity and cost that comes with that. NATS Streaming attempts to keep NATS’ spirit of simplicity, but it’s yet to be seen how it will reconcile that with the complex nature of distributed systems. I’m excited to see where Apcera takes NATS Streaming and generally the NATS ecosystem in the future since the team has a lot of experience in this area.

Stream Processing and Probabilistic Methods: Data at Scale

Stream processing and related abstractions have become all the rage following the rise of systems like Apache Kafka, Samza, and the Lambda architecture. Applying the idea of immutable, append-only event sourcing means we’re storing more data than ever before. However, as the cost of storage continues to decline, it’s becoming more feasible to store more data for longer periods of time. With immutability, how the data lives isn’t interesting anymore. It’s all about how it moves.

The shifting landscape of data architecture parallels the way we’re designing systems today. Specifically, the evolution of monolithic to service-oriented architecture necessitates a change in the way we do data integration. The traditional normalization approach doesn’t cut it. Our systems are composed of databases, caches, search indexes, data warehouses, and a multitude of other components. Moreover, there’s an increasing demand for online, real-time processing of this data that’s tantamount to the growing popularity of large-scale, offline processing along the lines of Hadoop. This presents an interesting set of new challenges, namely, how do we drink from the firehose without getting drenched?

The answer most likely lies in frameworks like Samza, Storm, and Spark Streaming. Similarly, tools like StatsD solve the problem of collecting real-time analytics. However, this discussion is meant to explore some of the primitives used in stream processing. The ideas extend far beyond event sourcing, generalizing to any type of data stream, unbounded or not.

Batch vs. Streaming

With offline or batch processing, we often have some heuristics which provide insight into our data set, and we can typically afford multiple passes of the data. Without a strict time constraint, data structures are less important. We can store the entire data set on disk (or perhaps across a distributed file system) and process it in batches.

With streaming data, however, there’s a need for near real-time processing—collecting analytics, monitoring and alerting, updating search indexes and caches, etc. With web crawlers, we process a stream of URLs and documents and produce indexed content. With websites, we process a stream of page views and update various counters and gauges. With email, we process a stream of text and produce a filtered, spam-free inbox. These cases involve massive, often limitless data sets. Processing that data online can’t be done with the conventional ETL or MapReduce-style methods, and unless you’re doing windowed processing, it’s entirely impractical to store that data in memory.

Framing the Problem

As a concrete example of stream processing, imagine we want to count the number of distinct document views across a large corpus, say, Wikipedia. A naive solution would be to use a hash table which maps a document to a count. Wikipedia has roughly 35 million pages. Let’s assume each document is identified by a 16-byte GUID, and the counters are stored as 8-byte integers. This means we need in the ball park of a gigabyte of memory. Given today’s hardware, this might not sound completely unreasonable, but now let’s say we want to track views per unique IP address. We see that this approach quickly becomes intractable.

To illustrate further, consider how to count the cardinality of IP addresses which access our website. Instead of a hash table, we can use a bitmap or sparse bit array to count addresses. There are over four billion possible distinct IPv4 addresses. Sure, we could allocate half a gigabyte of RAM, but if you’re developing performance-critical systems, large heap sizes are going to kill you, and this overhead doesn’t help.1

A Probabilistic Approach

Instead, we turn to probabilistic ways of solving these problems. Probabilistic algorithms and data structures seem to be oft-overlooked, or maybe purposely ignored, by system designers despite the theory behind them having been around for a long time in many cases. The goal should be to move these from the world of academia to industry because they are immensely useful and widely neglected.

Probabilistic solutions trade space and performance for accuracy. While a loss in precision may be a cause for concern to some, the fact is with large data streams, we have to trade something off to get real-time insight. Much like the CAP theorem, we must choose between consistency (accuracy) and availability (online). With CAP, we can typically adjust the level in which that trade-off occurs. This space/performance-accuracy exchange behaves very much the same way.

The literature can get pretty dense, so let’s look at what some of these approaches are and the problems they solve in a way that’s (hopefully) understandable.

Bloom Filters

The Bloom filter is probably the most well-known and, conceptually, simplest probabilistic data structure. It also serves as a good foundation because there are a lot of twists you can put on it. The theory provides a kernel from which many other probabilistic solutions are derived, as we will see.

Bloom filters answer a simple question: is this element a member of a set? A normal set requires linear space because it stores every element. A Bloom filter doesn’t store the actual elements, it merely stores the “membership” of them. It uses sub-linear space opening the possibility for false positives, meaning there’s a non-zero probability it reports an item is in the set when it’s actually not. This has a wide range of applications. For example, a Bloom filter can be placed in front of a database. When a query for a piece of data comes in and the filter doesn’t contain it, we completely bypass the database.


The Bloom filter consists of a bit array of length and hash functions. Both of these parameters are configurable, but we can optimize them based on a desired rate of false positives. Each bit in the array is initially unset. When an element is “added” to the filter, it’s hashed by each of the functions, h1…hk, and modded by m, resulting in indices into the bit array. Each bit at the respective index is set.


To query the membership of an element, we hash it and check if each bit is set. If any of them are zero, we know the item isn’t in the set. What happens when two different elements hash to the same index? This is where false positives are introduced. If the bits aren’t set, we know the element wasn’t added, but if they are, it’s possible that some other element or group of elements hashed to the same indices. Bloom filters have false positives, but false negatives are impossible—a member will never be reported incorrectly as a non-member. Unfortunately, this also means items can’t be removed from the filter.

It’s clear that the likelihood of false positives increases with each element added to the filter. We can target a specific probability of false positives by selecting an optimal value for m and k for up to n insertions.2 While this implementation works well in practice, it has a drawback in that some elements are potentially more sensitive to false positives than others. We can solve this problem by partitioning the m bits among the k hash functions such that each one produces an index over its respective partition. As a result, each element is always described by exactly bits. This prevents any one element from being especially sensitive to false positives. Since calculating k hashes for every element is computationally expensive, we can actually perform a single hash and derive k hashes from it for a significant speedup.3

A Bloom filter eventually reaches a point where all bits are set, which means every query will indicate membership, effectively making the probability of false positives one. The problem with this is it requires a priori knowledge of the data set in order to select optimal parameters and avoid “overfilling.” Consequently, Bloom filters are ideal for offline processing and not so great for dealing with streams. Next, we’ll look at some variations of the Bloom filter which attempt to deal with this issue.

Scalable Bloom Filter

As we saw earlier, traditional Bloom filters are a great way to deal with set-membership problems in a space-efficient way, but they require knowing the size of the data set ahead of time in order to be effective. The Scalable Bloom Filter (SBF) was introduced by Almeida et al. as a way to cope with the capacity dilemma.

The Scalable Bloom Filter dynamically adapts to the size of the data set while enforcing a tight upper bound on the rate of false positives. Like the classic Bloom filter, false negatives are impossible. The SBF is essentially an array of Bloom filters with geometrically decreasing false-positive rates. New elements are added to the last filter. When this filter becomes “full”—more specifically, when it reaches a target fill ratio—a new filter is added with a tightened error probability. A tightening ratio, r, controls the growth of new filters.


Testing membership of an element consists of checking each of the filters. The geometrically decreasing false-positive rate of each filter is an interesting property. Since the fill ratio determines when a new filter is added, it turns out this progression can be modeled as a Taylor series which allows us to provide a tight upper bound on false positives using an optimal fill ratio of 0.5 (once the filter is 50% full, a new one is added). The compounded probability over the whole series converges to a target value, even accounting for an infinite series.

We can limit the error rate, but obviously the trade-off is that we end up allocating memory proportional to the size of the data set since we’re continuously adding filters. We also pay some computational cost on adds. Amortized, the cost of filter insertions is negligible, but for every add we must compute the current fill ratio of the filter to determine if a new filter needs to be added. Fortunately, we can optimize this by computing an estimated fill ratio. If we keep a running count of the items added to the filter, n, the approximate current fill ratio, p, can be obtained from the Taylor series expansion with p \approx 1-e^{-n/m} where m is the number of bits in the filter. Calculating this estimate turns out to be quite a bit faster than computing the actual fill ratio—fewer memory reads.

To provide some context around performance, adds in my Go implementation of a classic Bloom filter take about 166 ns on my MacBook Pro, and membership tests take 118 ns. With the SBF, adds take 422 ns and tests 113 ns. This isn’t a completely fair comparison since, as the SBF grows over time, tests require scanning through each filter, but certainly the add numbers are intuitive.

Scalable Bloom Filters are useful for cases where the size of the data set isn’t known a priori and memory constraints aren’t of particular concern. It’s an effective variation of a regular Bloom filter which generally requires space allocation orders-of-magnitude larger than the data set to allow enough headroom.

Stable Bloom Filter

Classic Bloom filters aren’t great for streaming data, and Scalable Bloom Filters, while a better option, still present a memory problem. Another derivative, the Stable Bloom Filter (SBF), was proposed by Deng and Rafiei as a technique for detecting duplicates in unbounded data streams with limited space. Most solutions work by dividing the stream into fixed-size windows and solving the problem within that discrete space. This allows the use of traditional methods like the Bloom filter.

The SBF attempts to approximate duplicates without the use of windows. Clearly, we can’t store the entire history of an infinite stream in limited memory. Instead, the thinking is more recent data has more value than stale data in many scenarios. As an example, web crawlers may not care about redundantly fetching a web page that was crawled a long time ago compared to fetching a page that was crawled more recently. With this impetus, the SBF works by representing more recent data while discarding old information.

The Stable Bloom Filter tweaks the classic Bloom filter by replacing the bit array with an array of m cells, each allocated d bits. The cells are simply counters, initially set to zero. The maximum value, Max, of a cell is 2^d-1. When an element is added, we first have to ensure space for it. This is done by selecting P random cells and decrementing their counters by one, clamping to zero. Next, we hash the element to k cells and set their counters to Max. Testing membership consists of probing the k cells. If any of them are zero, it’s not a member. The Bloom filter is actually a special case of an SBF where d is one and P is zero.


Like traditional Bloom filters, an SBF has a non-zero probability of false positives, which is controlled by several parameters. Unlike the Bloom filter, an SBF has a tight upper bound on the rate of false positives while introducing a non-zero rate of false negatives. The false-positive rate of a classic Bloom filter eventually reaches one, after which all queries result in a false positive. The stable-point property of an SBF means the false-positive rate asymptotically approaches a configurable fixed constant.

Stable Bloom Filters lend themselves to situations where the size of the data set isn’t known ahead of time and memory is bounded. For example, an SBF can be used to deduplicate events from an unbounded event stream with a specified upper bound on false positives and minimal false negatives. In particular, if the stream is not uniformly distributed, meaning duplicates are likely to be grouped closer together, the rate of false positives becomes immaterial.

Counting Bloom Filter

Conventional Bloom filters don’t allow items to be removed. The Stable Bloom Filter, on the other hand, works by evicting old data, but its API doesn’t expose a way to remove specific elements. The Counting Bloom Filter (CBF) is yet another twist on this structure. Introduced by Fan et al. in Summary Cache: A Scalable Wide-Area Web Cache Sharing Protocol, the CBF provides a way of deleting data from a filter.

It’s actually a very straightforward approach. The filter comprises an array of n-bit buckets similar to the Stable Bloom Filter. When an item is added, the corresponding counters are incremented, and when it’s removed, the counters are decremented.

Obviously, a CBF takes n-times more space than a regular Bloom filter, but it also has a scalability limit. Unless items are removed frequently enough, a counting filter’s false-positive probability will approach one. We need to dimension it accordingly, and this normally requires inferring from our data set. Likewise, since the CBF allows removals, it exposes an opportunity for false negatives. The multi-bit counters diminish the rate, but it’s important to be cognizant of this property so the CBF can be applied appropriately.

Inverse Bloom Filter

One of the distinguishing features of Bloom filters is the guarantee of no false negatives. This can be a valuable invariant, but what if we want the opposite of that? Is there an efficient data structure that can offer us no false positives with the possibility of false negatives? The answer is deceptively simple.

Jeff Hodges describes a rather trivial—yet strikingly pragmatic—approach which he calls the “opposite of a Bloom filter.” Since that name is a bit of a mouthful, I’m referring to this as the Inverse Bloom Filter (IBF).

The Inverse Bloom Filter may report a false negative but can never report a false positive. That is, it may indicate that an item has not been seen when it actually has, but it will never report an item as seen which it hasn’t come across. The IBF behaves in a similar manner to a fixed-size hash map of m buckets which doesn’t handle conflicts, but it provides lock-free concurrency using an underlying CAS.

The test-and-add operation hashes an element to an index in the filter. We take the value at that index while atomically storing the new value, then the old value is compared to the new one. If they differ, the new value wasn’t a member. If the values are equivalent, it was in the set.


The Inverse Bloom Filter is a nice option for dealing with unbounded streams or large data sets due to its limited memory usage. If duplicates are close together, the rate of false negatives becomes vanishingly small with an adequately sized filter.


Let’s revisit the problem of counting distinct IP addresses visiting a website. One solution might be to allocate a sparse bit array proportional to the number of IPv4 addresses. Alternatively, we could combine a Bloom filter with a counter. When an address comes in and it’s not in the filter, increment the count. While these both work, they’re not likely to scale very well. Rather, we can apply a probabilistic data structure known as the HyperLogLog (HLL). First presented by Flajolet et al. in 2007, HyperLogLog is an algorithm which approximately counts the number of distinct elements, or cardinality, of a multiset (a set which allows multiple occurrences of its elements).

Imagine our stream as a series of coin tosses. Someone tells you the most heads they flipped in a row was three. You can guess that they didn’t flip the coin very many times. However, if they flipped 20 heads in a row, it probably took them quite a while. This seems like a really unconvincing way of estimating the number of tosses, but from this kernel of thought grows a fruitful idea.

HLL replaces heads and tails with zeros and ones. Instead of counting the longest run of heads, it counts the longest run of leading zeros in a binary hash. Using a good hash function means that the values are, more or less, statistically independent and uniformly distributed. If the maximum run of leading zeros is n, the number of distinct elements in the set is approximately 2^n. But wait, what’s the math behind this? If we think about it, half of all binary numbers start with 1. Each additional bit divides the probability of a run further in half; that is, 25% start with 01, 12.5% start with 001, 6.25% start with 0001, ad infinitum. Thus, the probability of a run of length n is 2^{-(n+1)}.

Like flipping a coin, there’s potential for an enormous amount of variance with this technique. For instance, it’s entirely possible—though unlikely—that you flip 20 heads in a row on your first try. One experiment doesn’t provide enough data, so instead you flip 10 coins. HyperLogLog uses a similar strategy to reduce variance by splitting the stream up across a set of buckets, or registers, and applying the counting algorithm on the values in each one. It tracks the longest run of zeros in every register, and the total cardinality is computed by taking the harmonic mean across all registers. The harmonic mean is used to discount outliers since the distribution of values tends to skew towards the right. We can increase accuracy by adding more registers, which of course comes at the expense of performance and memory.

HyperLogLog is a remarkable algorithm. It almost feels like magic, and it’s exceptionally useful for working with data streams. HLL has a fraction of the memory footprint of other solutions. In fact, it’s usually several orders of magnitude while remaining surprisingly accurate.

Count-Min Sketch

HyperLogLog is a great option to efficiently count the number of distinct elements in a stream using a minimal amount of space, but it only gives us cardinality. What about counting the frequencies of specific elements? Cormode and Muthukrishnan developed the Count-Min sketch (CM sketch) to approximate the occurrences of different event types from a stream of events.

In contrast to a hash table, the Count-Min sketch uses sub-linear space to count frequencies. It consists of a matrix with w columns and d rows. These parameters determine the trade-off between space/time constraints and accuracy. Each row has an associated hash function. When an element arrives, it’s hashed for each row. The corresponding index in the rows are incremented by one. In this regard, the CM sketch shares some similarities with the Bloom filter.


The frequency of an element is estimated by taking the minimum of all the element’s respective counter values. The thought process here is that there is possibility for collisions between elements, which affects the counters for multiple items.  Taking the minimum count results in a closer approximation.

The Count-Min sketch is an excellent tool for counting problems for the same reasons as HyperLogLog. It has a lot of similarities to Bloom filters, but where Bloom filters effectively represent sets, the CM sketch considers multisets.


The last probabilistic technique we’ll briefly look at is MinHash. This algorithm, invented by Andrei Broder, is used to quickly estimate the similarity between two sets. This has a variety of uses, such as detecting duplicate bodies of text and clustering or comparing documents.

MinHash works, in part, by using the Jaccard coefficient, a statistic which represents the size of the intersection of two sets divided by the size of the union:

J(A,B) = {{|A \cap B|}\over{|A \cup B|}}

This provides a measure of how similar the two sets are, but computing the intersection and union is expensive. Instead, MinHash essentially works by comparing randomly selected subsets through element hashing. It resembles locality-sensitive hashing (LSH), which means that similar items have a greater tendency to map to the same hash values. This varies from most hashing schemes, which attempt to avoid collisions. Instead, LSH is designed to maximize collisions of similar elements.

MinHash is one of the more difficult algorithms to fully grasp, and I can’t even begin to provide a proper explanation. If you’re interested in how it works, read the literature but know that there are a number of variations of it. The important thing is to know it exists and what problems it can be applied to.

In Practice

We’ve gone over several fundamental primitives for processing large data sets and streams while solving a few different types of problems. Bloom filters can be used to reduce disk reads and other expensive operations. They can also be purposed to detect duplicate events in a stream or prune a large decision tree. HyperLogLog excels at counting the number of distinct items in a stream while using minimal space, and the Count-Min sketch tracks the frequency of particular elements. Lastly, the MinHash method provides an efficient mechanism for comparing the similarity between documents. This has a number of applications, namely around identifying duplicates and characterizing content.

While there are no doubt countless implementations for most of these data structures and algorithms, I’ve been putting together a Go library geared towards providing efficient techniques for stream processing called Boom Filters. It includes implementations for all of the probabilistic data structures outlined above. As streaming data and real-time consumption grow more and more prominent, it will become important that we have the tools and understanding to deal with them. If nothing else, it’s valuable to be aware of probabilistic methods because they often provide accurate results at a fraction of the cost. These things aren’t just academic research, they form the basis for building online, high-performance systems at scale.

  1. Granted, if you’re sensitive to garbage-collection pauses, you may be better off using something like C or Rust, but that’s not always the most practical option. []
  2. The math behind determining optimal m and k is left as an exercise for the reader. []
  3. Less Hashing, Same Performance: Building a Better Bloom Filter discusses the use of two hash functions to simulate additional hashes. We can use a 64-bit hash, such as FNV, and use the upper and lower 32-bits as two different hashes. []