Recently I’ve been analyzing various data structures and algorithms in relaxed-memory (aka weak-memory) concurrency. The plan is to identify synchronization patterns that is used again and again, and explain concurrent programs in terms of the discovered synchronization patterns. This post reports the observations I’ve made so far. Most notably, and quite surprisingly, it seems most concurrent data structures and algorithms can be explained with only three synchronization patterns!

Why I am doing it? Because I badly needed it when I started learning concurrent programming. I read source codes of many data structures and articles that explain why they are correct. Each data structure made sense, in very subtle and yet intriguing reasons, but a question always remained: “how could they think of this data structure?” It becomes even worse when relaxed-memory concurrency comes to play: most articles on shared-memory concurrency, from blog posts to published papers, ignore the relaxed behaviors due to hardware and compiler optimizations, saying only that fences should be somehow properly inserted to implementation for restricting the relaxed behaviors. In short, at least to me, the whole concurrency business just looked like a black magic. I hoped some light be shed on it.

Before delving into the synchronization patterns I’ve found, as a background study, I’ll first review shared-memory concurrency and its relaxed behaviors. (For experts: I will deviate from the conventional explanation based on the “happens-before” relation, for the reasons I will explain later.)

Shared-Memory Concurrency and Relaxed Behaviors

Let’s start with a single-threaded program. A memory basically looks like a map from addresses to values:

memory : Map < Address , Value >

This should be the case regardless of whether there are caches in the memory hierarchy, or whether CPUs and compilers reorder load and store instructions to the memory. In fact, all these “optimizations” should guarantee that the memory appears to be a map from addresses to values for single-threaded programs.

So far so good. Now let’s think of a multi-threaded program. There are multiple threads that share a single memory so that a value stored by a thread can be loaded by another one. Unfortunately, due to all the hardware and compiler optimizations introduced for single-threaded programs, multi-threaded programs cannot enjoy the luxury of fantasy that the memory is just a map. To see this, let’s consider the following program:

fn main () { let a : AtomicUsize = 0 ; let b : AtomicUsize = 0 ; let th1 = fork (|| { a .store ( 42 , Relaxed ); b .load ( Relaxed ) }); let th2 = fork (|| { b .store ( 42 , Relaxed ); a .load ( Relaxed ) }); assert ( th1 .join () == 42 || th2 .join () == 42 ); }

Let’s say it’s written in an imaginary language whose syntax resembles that of Rust. a and b are unsigned integers ( AtomicUsize ) which are initially 0 , a forked thread th1 stores 42 to a and then loads from b , and another forked thread th2 stores 42 to b and then loads from a . Relaxed roughly means the loads and stores should be compiled to plain load and store instructions in the assembly. The question of the assertion is, is it always true that either the value loaded from a equals to 42, or that loaded from b equals to 42? At first glance, it seems trivially so, since 42 should be stored to either a or b before loading a value from either of them. However, it is not the case for mysterious reasons, and the assertion may fail by reading 0 from both a and b . The moral of this story is that the abstraction of memory as a map is broken, and programs are allowed to exhibit more behaviors, even in unexpected ways from the programmer’s point of view.

In order to tame these behaviors, a program should be properly synchronized. CPU architectures such as x86-64, ARMv7/8, POWER, and programming languages such as C, C++, LLVM, Rust support synchronization primitives in addition to plain load and store instructions. Unfortunately, these primitives had been really difficult to understand, and even more difficult to program with. As a result, only a small group of experts have spoken these primitives and constructed concurrency libraries for the others. Still, these experts often misuse the primitives, and even worse, the ISA and language standards have bugs in their specifications!

In order to resolve this issue, last year my collaborators and I developed a promising semantics for relaxed-memory concurrency that clearly explains the semantics of C/C++ synchronization primitives. Now I will explain some of its concepts, namely coherence and view, that collectively answer the very question of memory consistency models: which value(s) a thread can read from the memory? As you will see, these concepts will play a key role in explaining what is synchronization and how the synchronization patterns work.

Coherence and View

Let’s first see how a memory looks like. A memory in the promising semantics is a map from addresses to a set of values uniquely ordered by a timestamp:

memory : Map < Address , Map < Timestamp , Value >>

We call each value a message. When a thread stores a value to an address, a message with the stored value is added to the address; when a thread loads from an address, the value of a message of the address is returned. It has two implications: (1) threads may be able to concurrently read different values from the same address; but (2) at least, it is guaranteed that there is “some” global order (namely the timestamp) of the writes to the same address. We call this the “coherence order” (or “modification order” in the C/C++ standards lingo). For example, the messages of the address x in a memory may look like:

[x=1@10] [x=42@20] [x=37@30] [x=2@40] [x=3@50] x --------+---------+---------+---------+---------+-------

where there are 5 messages of the address x : value 1 at the timestamp 10, 42 at 20, 37@30, 2@40, and 3@50. It is worth noting that timestamps of different addresses are not related at all, since a coherence order is specific to a single address.

When loading from and storing to the memory, threads should respect the coherence order: once a thread acknowledged a message, it cannot read from and write to a timestamp before the message w.r.t. the coherence order. More precisely, each thread maintains a thread view, which is a map from addresses to timestamps that tracks the maximum timestamp it acknowledged for each address:

view : Map < Address , Timestamp >

And the thread view interacts with load and store instructions as follows. Suppose a thread’s view on the address x is at the timestamp t1 ( view[x] == t1 ). In other words, the thread somehow acknowledged the message of the address x at timestamp t1 , and nothing later. When it reads a message at timestamp t2 from x , t1 <= t2 should hold, and at the same time view[x] is updated to t2 (since the message at t2 is just acknowledged). When it writes a message at timestamp t2 to x , t1 < t2 should hold, and at the same time view[x] is updated to t2 (since the message at t2 is automatically acknowledged). In short, loads and stores update a thread’s view, and view should be non-decreasing.

This coherence condition is arguably a minimal requirement of sanity, so (almost) all CPUs and compilers guarantee this by default, even in the absence of any synchronization primitives. (For experts: the coherence condition of the promising semantics is roughly comparable to the coherence axioms in the C/C++ standards. See the paper for more detailed comparison.)

What is Synchronization?

Coherence is good, but with it alone programmers cannot construct concurrent programs. This is mainly because the coherence condition applies to a single address at a time: acknowledgment of a message of an address x has nothing to do with the thread’s view on another address y . It is synchronization that relates acknowledgment of multiple addresses.

I analyzed various real-world concurrent data structures and algorithms, and classified the synchronization patterns used in them into three categories: positive piggybacking synchronization, negative piggybacking synchronization, and interleaving synchronization. In the remaining of this post, I will explain these three patterns and explain a few concurrent data structures with them as running examples.

Let’s begin with the most basic and well-understood pattern: positive piggybacking synchronization. (For experts: think Release - Acquire synchronization. I deliberately used general terms rather than C/C++-specific terms, e.g. “positive” and “piggybacking” rather than “ Release - Acquire ”, for implying that these patterns apply to CPUs and other languages as well.)

Pattern 1: Positive Piggybacking Synchronization

In this pattern, we relate multiple addresses by piggybacking the knowledge on an address to another address. For example, consider the following program:

fn main () { let data : AtomicUsize = 0 ; let flag : AtomicUsize = 0 ; let th1 = fork (|| { data .store ( 42 , Relaxed ); flag .store ( 1 , Release ) }); let th2 = fork (|| { if ( flag .load ( Acquire )) { assert ( data .load ( Relaxed ) == 42 ); } }); }

Here, the thread th1 writes 42 to data , and then set the flag . The thread th2 checks if the flag is set, and if so, assert that data always contains 42. In order for successful assertion, knowledge on the data variable should be piggybacked on the flag variable.

This synchronization is the job of the Release and Acquire ordering annotations in the load and store instructions to flag . When a thread writes a message with the Release ordering, the thread’s view is annotated in the written message; and when a reads a message with the Acquire ordering, the message’s view is acknowledged by the reading thread. In order to do so, we attach a view to memory’s messages:

memory: Map<Address, Map<Timestamp, (Value, View)>>

Let’s see what happens for the example above. Suppose th1 stored data = 42 at timestamp 10 , and flag = 1 at timestamp 5, as described in the timeline below. At the time of writing flag = 1 , the thread th1 ’s view is [data@10, flag@5] , and this view is annotated in the message flag@5 because of the Release ordering. When th2 reads flag = 1 , its view becomes [data@10, flag@5] thanks to the Acquire ordering, forcing the later load from data to read 42 (or a message written in a later timestamp).

[data=42@10] data ---------+----------- [flag=1@5, view=[data@10, flag@5]] flag ----+---------------------------------

Example: Spinlock

This positive piggybacking synchronization pattern is used to implement a spinlock, which provides mutual exclusion (mutex) among the critical sections (or code regions) in multiple threads. By “mutual exclusion” I mean (1) there is a total order over critical sections; and (2) the view at the end of a critical section should be acknowledged at the beginning of a later critical section. (For experts: it is different from the conventional requirement that a critical section happens before a later one.)

The following is an implementation of spinlock. Function new() creates a new spinlock, lock() marks the beginning of a critical section, and unlock() marks the end of it:

struct Spinlock { lock : AtomicUsize , } impl Spinlock { fn new () -> Spinlock { Spinlock { lock : 0 , } } fn lock ( & self ) { while ( self .lock .compare_and_swap ( 0 , 1 , Acquire ) != 0 ) {} } fn unlock ( & self ) { self .lock .store ( 0 , Release ); } }

A spinlock consists of an integer ( AtomicUsize ) variable lock , which represents whether it is locked by a critical section or not. As described in the timeline below, if its last value w.r.t. the coherence order is 0 , it is not locked; if it is 1 , it is locked:

[UNLOCKED] (init) (lock)-(unlock) [0] [1] [0] lock +!!!!!!+------+------- [LOCKED] (init) (lock)-(unlock) (lock) [0] [1] [0] [1] lock +!!!!!!+------+!!!!!!!!!!+----- [UNLOCKED, again] (init) (lock)-(unlock) (lock)----(unlock) [0] [1] [0] [1] [0] lock +!!!!!!+------+!!!!!!!!!!+---------+-------

The new() function initializes a spinlock by assigning 0 to the internal lock variable, meaning that it is not locked at the beginning.

The lock() function spins until successfully updating the variable from 0 to 1 . In order to guarantee that only one thread can update a variable from 0 to 1 , the compare_and_swap() function resolves the race over the lock variable. In order for a thread to win the race and enter a critical section, the thread should be able to read the old value ( 0 in this case) and then subsequently write a new value ( 1 in this case), with no existing messages between the old and the new value. If it is the case, the timestamps between the old and the new value are marked as unusable in the future (represented as ! in the timeline), and the old value is returned. For example, a successful lock() may change the memory from [UNLOCKED] to [LOCKED] in the above timeline. Thanks to the exclusive nature of compare_and_swap() (and the invariant that the neighbor of only the last message is not marked as unusable), at any time, only one thread can successfully update the lock variable. If a thread loses the race, compare_and_swap() returns a value in the memory just like load() , letting the lock() function try again.

The unlock() function simply stores 0 to the lock variable. For example, unlock() may change the memory from [LOCKED] to [UNLOCKED, again] .

Now let’s see how the above implementation guarantees mutual exclusion. First, the view at the end of a critical section is annotated in the lock variable thanks to the Release ordering of unlock() ’s store to the lock variable. Then the annotated view is acknowledged at the beginning of the next critical section due to the Acquire ordering of lock() ’s compare_and_swap() . This (positive) piggybacking synchronization on the lock variable transitively sends the view of a critical section to later ones.

Variations

There are other kinds of piggybacking synchronization than Release - Acquire synchronization. In theory, you can think of the following dimensions of variation:

The bridge of piggybacking. In Release - Acquire synchronization, a Release -write to the piggybacking address ( flag in the example) and an Acquire -read from the address form a “bridge” of piggybacking. We may call this WR (write-read) bridge. You can imagine RRc (read-read via coherence) bridge, whose two reads from the piggybacking address are ordered by the coherence order; RWc bridge of a read and a coherence-later write to the piggybacking address; WRc bridge; and WWc bridge. (For experts: you may call C/C++ release sequence “(WR)* bridge” (“*” means the Kleene Star).)

Synchronization marker. In the example, the Release and Acquire orderings are annotated in the store and load instructions. Instead, you can mark a program point earlier than the bridge’s write as Release , and mark a program point later than the bridge’s load as Acquire ; we call these markers “fences”. For example, the example could have been written in this way: // ... unchanged let th1 = fork (|| { data .store ( 42 , Relaxed ); fence ( Release ); flag .store ( 1 , Relaxed ) }); let th2 = fork (|| { if ( flag .load ( Relaxed )) { fence ( Acquire ); assert ( data .load ( Relaxed ) == 42 ); } }); Here, the view at the time of fence(Release) is acknowledged at the time of fence(Acquire) , thereby successfully asserting data.load(Relaxed) == 42 . You may mix Release store with Acquire fence or Release fence with Acquire load for achieving the same thing.

In reality, not all these combinations are realized in CPUs and languages, partially because they are not efficiently implementable. For what is worth, C/C++ supports (1) synchronization via WR (and WRu) bridges annotated with Release and Acquire orderings, namely the ordinary Release - Acquire synchronization; (2) its variations with fences; and (3) synchronization via RRc, RWc, WRc, WWc bridges in the presence of SeqCst fences in both sides, where SeqCst is the strongest and the most expensive ordering in C/C++.

Note

Positive piggybacking synchronization is a relatively well-understood pattern compared to the other two. Many of the earliest concurrent data structures, e.g. Treiber stack and Michael-Scott queue, rely only on this pattern. In C/C++, the notion of “happens-before” relation, which is at the heart of the memory model, is built upon this pattern only. While there is no satisfactory program logic for relaxed-memory concurrency in general, a program logic for release/acquire fragments of C/C++ had been successfully developed.

This pattern is “positive” in the sense that the synchronization depends on positive information, namely the acknowledgment of a message. For the example program, since th2 observed flag = 1 , it should have also observed data = 42 . The next pattern is also based on piggybacking, but it is “negative”: it depends on negative information, namely the absence of acknowledgment.

Pattern 2: Negative Piggybacking Synchronization

For the example program above, acknowledgment of flag = 1 implies that of data = 42 . The contraposition is also true: the absence of the acknowledgment of data = 42 implies that of the acknowledgment of flag = 1 . This pattern is used in “sequence lock”: an optimized implementation of reader-writer lock. Note that a reader-writer lock is a mechanism for protecting data which guarantees mutual exclusion among writers, while providing a designated, optimized method for reading, but not writing, the protected data. For correctness, the reader should be atomic in the sense that it should not observe intermediate modification of data by a concurrent writer.

Example: Sequence Lock

The following is an implementation of sequence lock. Function new() creates a new sequence lock, writer_lock() and writer_unlock() marks the beginning and the end of a critical section in which the writer can access the protected data, and read() returns the protected data:

struct < T > Seqlock < T : Copy > { seq : AtomicUsize , data : Atomic < T > , } impl < T : Copy > Seqlock < T > { fn new ( data : T ) -> Seqlock < T > { Seqlock { seq : 0 , data : Atomic :: new ( data ), } } fn writer_lock ( & self ) -> ( usize , & mut T ) { loop { let seq = self .seq .load ( Relaxed ); if ( seq & 1 != 0 ) { continue }; if ( self .seq .compare_and_swap ( seq , seq + 1 , Acquire ) != seq ) { continue }; fence ( Release ); return ( seq + 2 , self .data as & mut T ); // It's not Rust exactly.. } } fn writer_unlock ( & self , seq : usize ) { self .seq .store ( seq , Release ); } fn read ( & self ) -> T { loop { let seq1 = self .seq .load ( Acquire ); if ( seq1 & 1 != 0 ) { continue }; let result = self .data .load ( Relaxed ); fence ( Acquire ); let seq2 = self .seq .load ( Relaxed ); if ( seq1 != seq2 ) { continue }; return result ; } } } // Using a seqlock fn main () { let seqlock = Seqlock :: new ( ... ); let th1 = fork (|| { let ( seq_next , val ) = seqlock .writer_lock (); ... // writer's critical section seqlock .writer_unlock ( seq_next ); }); let th2 = fork (|| { let val = seqlock .read (); }); }

The new() , writer_lock() , and writer_unlock() functions guarantee mutual exclusion for the same reason with spinlock, but using even numbers instead of 0 for representing unlocked states and odd numbers instead of 1 for representing locked states. Below is an example timeline for the seq variable. For example, a writer updates seq from 0 to 1 and then writes 2 to seq . Let’s call it W2 . Similarly, the writer that writes 4 to seq is W4 , and so on:

(R4) (init) (W2: lock)-(unlock) (W4: lock)----(unlock) (W6: lock)--(unlock) [0] [1] [2] [3] [4] [5] [6] seq +!!!!!!+----------+!!!!!!!!+-------------+!!!!!!!!!+-----------+-------

The read() function is atomic for the following reasons. Suppose a reader R4 observed seq1 = seq2 = 4 . I will show that R4 reads the data written by W4 .

First, the view at the end of W4 is sent to the beginning of R4 by positive piggybacking synchronization from writer_unlock() ’s Release -write ( self.seq.store(seq, Release) ) to read() ’s Acquire -load ( let seq1 = self.seq.load(Acquire) ). In particular, we know that the view on data at the end of W4 <= view on data at the beginning of R4 .

Second, the view on data at the end of R4 <= the view on data at the end of W4 . Otherwise, a part of the data R4 read from came from a writer later than W4 . By positive piggybacking synchronization on data from the later writer’s writer_lock() ’s fence(Release) to read() ’s fence(Acquire) , the message of seq = 5 or later should have been acknowledged after read() ’s fence(Acquire) . But it is a contradiction, because the reader observed that seq2 = 4 . In other words, by negative piggybacking synchronization on data, seq2 = 4 means the view on data at the end of R4 <= the view on data at the end of W4 .

Therefore throughout the execution of R4 , its view on data should equal to the view on data at the end of W4 . So R4 reads exactly the data completely written by W4 .

(For experts: it is worth noting that R4 does not happen before W6 : we only know that R4 ’s view on data <= the view on data at the beginning of W6 . It is just sufficient for a reader-writer lock to be correct. In my opinion, the specifications of some data structures, including sequence lock, are more naturally expressible with views than with the happens-before relation. It’s the reason I prefer explaining the synchronization patterns with views.)

Note

Recall that in the synchronization patterns introduced so far, knowledge on an address is piggybacked on another address, reusing the coherence order of the “another address” as a bridge. But in some cases, we need a stronger synchronization by ordering arbitrary program points from different threads. This is the job of interleaving fences.

Pattern 3: Interleaving Synchronization

Interleaving fences (the SeqCst fence in C/C++, and the most heavyweight fences in CPU architectures) mark the program points that should be totally ordered. When a thread th1 executes an interleaving fence before another thread th2 executes another interleaving fence w.r.t. the total order, th1 ’s view before executing the fence should be acknowledged by th2 after executing the fence. In order to do so, the promising semantics maintains a global view for interleaving fences:

static mut interleaving_view : View // we need `unsafe` accessor in Rust, but..

When a thread executes an interleaving fence, it calculates the (address-wise) max of its view and the global interleaving view, and set the max view to its own view and the global interleaving view:

fn execute_interleaving_fence ( & mut thread ) { let view = max ( thread .view , interleaving_view ); thread .view = interleaving_view = view ; }

Interleaving synchronization is very powerful: in fact, it subsumes both forms of piggybacking synchronizations. However, in my opinion, it is more difficult to reason about than piggybacking synchronizations, as combinatorial explosion occurs when analyzing all possible interleavings. I would recommend to use this pattern only if the power of interleaving is actually necessary.

This is the only kind of synchronization supported in what is called “sequentially consistent semantics”, or “interleaving semantics”, where all the instructions are interleaving by default (thereby allowing no relaxed behaviors). For this reason, I believe sequentially consistent semantics is difficult to reason about, at least as much as relaxed-memory concurrency semantics. I know many of you cannot agree with me on this matter; sequential consistency is regarded as the ideal and the easiest semantics for shared-memory concurrency for decades. But I believe the scene has changed a little bit: now we can explain the synchronization patterns in terms of views.

Interleaving synchronization is used, for example, in Peterson’s algorithm, which is an early mutual exclusion algorithm, and work-stealing deque by Chase and Lev. In the remaining of this section, I will analyze Peterson’s algorithm in more details.

Example: Peterson’s Mutual Exclusion

The following is an implementation of Peterson’s mutual exclusion algorithm. As opposed to spinlock and sequence lock, Peterson’s algorithm I am presenting here supports only two threads:

fn main () { let flag : [ AtomicBool ; 2 ]; let turn : AtomicUsize = 0 ; fn lock ( id : Usize ) { flag [ id ] .store ( true , Relaxed ); fence ( SeqCst ); // A turn .store ( 1 - id , Relaxed ); fence ( SeqCst ); // B while ( flag [ 1 - id ] .load ( Acquire ) && turn .load ( Relaxed ) == 1 - id ) {} } fn unlock ( id : Usize ) { flag [ id ] .store ( false , Release ); } let th0 = fork (|| { lock ( 0 ); // critical section unlock ( 0 ); }); let th1 = fork (|| { lock ( 1 ); // critical section unlock ( 1 ); }); }

Peterson’s algorithm guarantees mutual exclusion for the following reasons. In th0 ’s and th1 ’s call to lock() , there are four SeqCst fences: th0 ’s first fence ( A0 ), th0 ’s second fence(SeqCst) ( B0 ), th1 ’s first fence ( A1 ), and th1 ’s second fence ( B1 ). We will analyze every possible order of these fences, but without loss of generality, it is sufficient to analyze the following orders only:

A0 -> B0 -> A1 -> B1 . By interleaving property, flag[0] = true and turn = 1 should be acknowledged after A1 . So th1 should write turn = 0 after turn = 1 w.r.t. the coherence order, and it should spin until th0 write flag[0] = false in unlock() . By positive piggybacking synchronization on flag[0] , the view at the end of th0 ’s critical section is sent to the beginning of th1 s critical section.

A0 -> A1 -> B0 -> B1 or A0 -> A1 -> B1 -> B0 . By interleaving property, flag[0] = true and flag[1] = true should be acknowledged after both B0 and B1 . Without loss of generality, suppose that th0 wrote turn = 1 before th1 did turn = 0 w.r.t. the coherence order. Then th1 ’s lock() should spin until th0 writes flag[0] = false in unlock() . By positive piggybacking synchronization on flag[0] , the view at the end of th0 ’s critical section is sent to the beginning of th1 s critical section.

Case Study: Crossbeam

So far I identified three relaxed-memory concurrency synchronization patterns, and explained a few data structures with the mix of those three patterns. Now let’s analyze something bigger: the Crossbeam library for epoch-based concurrent memory reclamation scheme in Rust. For more information on this library, I refer you to Aaron Turon’s introduction to Crossbeam. I wrote down a Crossbeam RFC that explains why Crossbeam’s implementation is correct w.r.t. the C/C++ memory model. (In fact, many of the ideas I am presenting here was conceived when I was writing the RFC.) I believe now you can read it, and find where, how, and which patterns are used in Crossbeam

Future Work

I tried to identify synchronization patterns as comprehensively as possible, but probably there should be more patterns yet to be discovered. Most notably, I omitted some emerging synchronization primitives, which have a potential to be a basis for new and faster synchronization patterns:

Data-dependent piggybacking synchronization. For example, Consume loads in C/C++ is a variant of Acquire loads where the effect of Acquire (acknowledging the Release d view) takes place only for the instructions that depend on the Consume -loaded value. For example, consider the following program: fn main () { let data : AtomicUsize = 0 ; let ptr : AtomicUsize = 0 ; let th1 = fork (|| { data .store ( 42 , Relaxed ); ptr .store ( & data as usize , Release ) }); let th2 = fork (|| { let p = ptr .load ( Consume ) as & AtomicUsize ; if ( ! p .is_null ()) { assert ( p .load ( Relaxed ) == 42 ); // dependent on `p`, safe assert ( data .load ( Relaxed ) == 42 ); // independent from `p`, unsafe } }); } Here, ptr is read with the Consume annotation. Since the assertion p.load(Relaxed) == 42 depends on the value p read from ptr , the piggybacking synchronization happens and the assertion should succeed. On the other hand, since the assertion data.load(Relaxed) == 42 does not syntactically depend on p , piggybacking synchronization does not happen, and data.load(Relaxed) can read the initial value 0 and fail the assertion. Consume or READ_ONCE is faster than Acquire in relaxed CPU architectures such as ARM and Power, and is actually used in the Linux kernel (in the name of READ_ONCE ). Unfortunately, we currently do not know of a good semantics for that. However, I believe its usage falls into the positive/negative piggybacking synchronization patterns.

Strongly-synchronizing load/store instructions. C/C++ allows SeqCst annotations for load and store instructions. These instructions are intended to be more strongly synchronizing than those annotated with just Release and Acquire . Even, in order to support more efficient compilation of them, the ARMv8 architecture introduced the LDA (load-acquire) and STL (store-release) instructions (and their variants). However, as discussed in this paper, the semantics of SeqCst loads and stores have been severely broken, and the only fix proposed so far is too complicated. Fixing its semantics and identifying its usage patterns is an important future work.

System-wide synchronization. The sys_membarrier syscall on Linux and FlushProcessWriteBuffers on Windows essentially perform a SeqCst fence on all CPU cores. As discussed in a Crossbeam RFC, these system-wide fences may be used to eliminate a fence in a critical path at the expense of introducing a system-wide fence in a cold path. Identifying its usage patterns is also an important future work.

Synchronizing heterogeneous systems. So far I focused only on on-board inter-CPU synchronization, but these days more hardware devices, including GPU, NIC (Network Interface Controller) cards, and maybe TPU (Tensorflow Processing Unit) are synchronizing with each other. Is there any unique usage pattern in these heterogeneous systems?

Conclusion

I hope this post got its job done: clarifying the essence of relaxed-memory concurrency by identifying synchronization patterns, thereby helping you start writing concurrent programs. I hope concurrency is no longer a black magic to you as was it to me two years ago, but a disciplined subject of modern systems programming. Happy hacking concurrency!

Edit

I would like to thank @foollbar, @stjepang, @Vtec234, Benjamin Fry, Derek Dreyer, and Gil Hur for their helpful comments to an earlier version of this post.