A deep dive into Multicore OCaml garbage collector

I recently gave a talk on the internals of multicore OCaml GC at Jane Street offices in NYC. The slides from the talk are available online. But I felt that the slides alone aren’t particularly edifying. This post is basically the slides from the talk annotated with notes.

Abstract

In a mostly functional language like OCaml, it is desirable to have each domain (our unit of parallelism) collect its own local garbage independently. Given that OCaml is commonly used for writing latency sensitive code such as trading systems, UIs, networked unikernels, it is also desirable to minimise the stop-the-world phases in the GC. Although obvious, the difficulty is to make this work in the presence of mutations and concurrency. In this talk, I will present the overall design of Multicore OCaml GC, but also deep dive into a few of the interesting techniques that make it work.

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Multicore OCaml project is led from OCaml Labs within the University of Cambridge. Multicore OCaml project is led from OCaml Labs within the University of Cambridge. Stephen Dolan begun the project while procrastinating on writing up his dissertation . Hurray to that!

Multicore OCaml extends OCaml with native support for concurrency and parallelism. Unlike many other languages, we clearly separate concurrency from parallelism in the language. Concurrency in Multicore OCaml is expressed through lightweight language level threads called fibers. The unit of parallelism is a domain, which maps to kernel threads. Many kernel threads may service a particular domain, but only one of those kernel threads ever runs OCaml code. A typical program is expected to have a large number of fibers mapped over a few domains. In this talk, I provide an overview of the Multicore OCaml GC design, with a few deep dives into some of the interesting and novel techniques. Multicore OCaml extends OCaml with native support for concurrency and parallelism. Unlike many other languages, we clearly separate concurrency from parallelism in the language. Concurrency in Multicore OCaml is expressed through lightweight language level threads called fibers. The unit of parallelism is a domain, which maps to kernel threads. Many kernel threads may service a particular domain, but only one of those kernel threads ever runs OCaml code. A typical program is expected to have a large number of fibers mapped over a few domains. In this talk, I provide an overview of the Multicore OCaml GC design, with a few deep dives into some of the interesting and novel techniques.

It is difficult to appreciate the subtleties of the choice of GC techniques in isolation. So we shall begin with a sane setup - a GC for a sequential purely functional language. We shall subsequently extend the language with series of reprehensible extensions including mutations, parallelism and concurrency, in that order, and observe how the sane setup falls apart and what we shall do to recover sanity while retaining efficiency. The early parts of the talk should be unsurprising to someone familiar with GC internals, but is useful for setting up the latter material. It is difficult to appreciate the subtleties of the choice of GC techniques in isolation. So we shall begin with a sane setup - a GC for a sequential purely functional language. We shall subsequently extend the language with series of reprehensible extensions including mutations, parallelism and concurrency, in that order, and observe how the sane setup falls apart and what we shall do to recover sanity while retaining efficiency. The early parts of the talk should be unsurprising to someone familiar with GC internals, but is useful for setting up the latter material.

The figure shows the snapshot of the runtime state of a sequential purely functional language. We have a set of objects allocated on the heap. These objects may be pointed to be other heap objects as well as any of the registers and the current runtime stack. The simplest GC strategy is to stop the program and perform a mark-and-sweep garbage collection. The core idea is that we start from the roots of the program state i.e., current stack and registers, and perform a depth-first traversal through the object graph. Any unreachable objects are garbage and they can be reclaimed. Tri-colour marking is a standard marking algorithm. During marking, the objects are in one of three states: white(unmarked), grey(marking) and black(marked). The figure shows the snapshot of the runtime state of a sequential purely functional language. We have a set of objects allocated on the heap. These objects may be pointed to be other heap objects as well as any of the registers and the current runtime stack. The simplest GC strategy is to stop the program and perform a mark-and-sweep garbage collection. The core idea is that we start from the roots of the program state i.e., current stack and registers, and perform a depth-first traversal through the object graph. Any unreachable objects are garbage and they can be reclaimed. Tri-colour marking is a standard marking algorithm. During marking, the objects are in one of three states: white(unmarked), grey(marking) and black(marked).

At the beginning of the GC, all objects are white. The GC marks grey any white object it finds and pushes it into the mark stack. Subsequently, objects are popped off the mark stack, all of its white children marked, and the object is marked black. We have the invariant that a black object does not point to a white object. This is called the tri-colour invariant. The figure shows the state when object A and its children have been marked (hence, A is black), object B has been marked grey and is on the mark stack. At the beginning of the GC, all objects are white. The GC marks grey any white object it finds and pushes it into the mark stack. Subsequently, objects are popped off the mark stack, all of its white children marked, and the object is marked black. We have the invariant that a black object does not point to a white object. This is called the tri-colour invariant. The figure shows the state when objectand its children have been marked (hence,is black), objecthas been marked grey and is on the mark stack.

GC is done when the mark stack is empty. At this point, all reachable objects are black. Any unreachable objects are white. A sweeper examines all the allocated objects, and marks any object still white as free space that can subsequently be reused. GC is done when the mark stack is empty. At this point, all reachable objects are black. Any unreachable objects are white. A sweeper examines all the allocated objects, and marks any object still white as free space that can subsequently be reused.

Mark and sweep GC has several advantages. For starters, it is a very simple algorithm as GC algorithms go. The algorithm is also naturally incremental. The mutator (OCaml code) and the GC work can alternate between each other, minimizing the pause times in the GC. However, the primary disadvantage with this scheme is that one needs to maintain a free-list of objects for allocations. While there are many algorithms to find the best place to allocate an object, all of them have non-trivial overheads. Functional programming languages have high rate of allocation and would benefit fast allocations. Moreover, free-list implementations also suffer from fragmentation. Mark and sweep GC has several advantages. For starters, it is a very simple algorithm as GC algorithms go. The algorithm is also naturally incremental. The mutator (OCaml code) and the GC work can alternate between each other, minimizing the pause times in the GC. However, the primary disadvantage with this scheme is that one needs to maintain a free-list of objects for allocations. While there are many algorithms to find the best place to allocate an object, all of them have non-trivial overheads. Functional programming languages have high rate of allocation and would benefit fast allocations. Moreover, free-list implementations also suffer from fragmentation.

To alleviate this, functional programming languages typically implement a generational GC. Generational hypothesis says that the young objects are far more likely to die than older objects. To take advantage of this, the heap is split into two -- a small minor heap, where new objects are allocated and a larger major heap. In particular, objects in the minor heap are allocated by bumping the frontier, leading to fast allocations. When the minor heap is full, we garbage collect the minor heap. To alleviate this, functional programming languages typically implement a generational GC. Generational hypothesis says that the young objects are far more likely to die than older objects. To take advantage of this, the heap is split into two -- a small minor heap, where new objects are allocated and a larger major heap. In particular, objects in the minor heap are allocated by bumping the frontier, leading to fast allocations. When the minor heap is full, we garbage collect the minor heap.

Minor heap is GCed with a copying collector, and any object that survives the minor GC is promoted to the major GC. A nice aspect of copying collection for minor GC is that only the live objects need to be scanned, unlike mark and sweep collection where the sweeper needs to examine every allocated object. Given the generational hypothesis, we will not examine most of the objects in the minor heap. As a data point, the minor GC survival rate while compiling the OCaml compiler is around 10%. Minor heap is GCed with a copying collector, and any object that survives the minor GC is promoted to the major GC. A nice aspect of copying collection for minor GC is that only the live objects need to be scanned, unlike mark and sweep collection where the sweeper needs to examine every allocated object. Given the generational hypothesis, we will not examine most of the objects in the minor heap. As a data point, the minor GC survival rate while compiling the OCaml compiler is around 10%.

The roots for the minor collection are the current stack and the live register set. Since our language is purely functional, there will be no pointers from the major heap to the minor heap. This is because all the objects in major heap are older than all the objects in the minor heap. The lack of mutations mean that an object can only point to an older object. Hence, we don't need to care about objects in the major heap for minor collections. The roots for the minor collection are the current stack and the live register set. Since our language is purely functional, there will be no pointers from the major heap to the minor heap. This is because all the objects in major heap are older than all the objects in the minor heap. The lack of mutations mean that an object can only point to an older object. Hence, we don't need to care about objects in the major heap for minor collections.

With mutations, the major heap may point to the minor heap. For example, by assigning a minor heap object to a major heap reference. We need to know these references so that we can treat them as roots for the minor collection. With mutations, the major heap may point to the minor heap. For example, by assigning a minor heap object to a major heap reference. We need to know these references so that we can treat them as roots for the minor collection.

Naively, we can scan the entire major heap for every minor GC to find such pointers. But this is impractical. Naively, we can scan the entire major heap for every minor GC to find such pointers. But this is impractical.

Instead, we intercept the writes with a write barrier that records such pointers in an auxiliary data structure called the remembered set. Instead, we intercept the writes with a write barrier that records such pointers in an auxiliary data structure called the remembered set.

Remembered set holds the set of pointers from the major heap to the minor heap, and is used as the root for minor collections. After the minor collection, the remembered set is cleared. Remembered set holds the set of pointers from the major heap to the minor heap, and is used as the root for minor collections. After the minor collection, the remembered set is cleared.

Mutations breaks tri-colour invariant. Suppose our heap has these three objects... Mutations breaks tri-colour invariant. Suppose our heap has these three objects...

and we assign the white object B to the black object C ... and we assign the white objectto the black object...

and we subsequently delete the pointer from A to B . If we perform a major GC now, and we subsequently delete the pointer fromto. If we perform a major GC now,

A will be eventually marked as black. But the white object B is only pointed to by the black object C , which will not be marked, will be eventually marked as black. But the white objectis only pointed to by the black object, which will not be marked,

and the sweeper will mark it as free. This leaves the heap in an inconsistent state. and the sweeper will mark it as free. This leaves the heap in an inconsistent state.

Mutations are problematic if two conditions hold. 1. there exists a black to white pointer and 2. all references from a grey object through a chain of white objects to that white object is deleted. We can recover correctness if we disallow one of these conditions. Mutations are problematic if two conditions hold. 1. there exists a black to white pointer and 2. all references from a grey object through a chain of white objects to that white object is deleted. We can recover correctness if we disallow one of these conditions.

An insertion barrier prohibits 1. Whenever a black to white pointer may be established, the insertion barrier marks the target. An insertion barrier prohibits 1. Whenever a black to white pointer may be established, the insertion barrier marks the target.

This prevents B from being GCed. The insertion barrier is said to preserve strong tri-colour invariant: A black object never points to a white object. This preventsfrom being GCed. The insertion barrier is said to preserve strong tri-colour invariant: A black object never points to a white object.

A deletion barrier prohibits 2. In particular, deletion barrier allows the black to white pointer from C to B . A deletion barrier prohibits 2. In particular, deletion barrier allows the black to white pointer fromto

But when the pointer from the grey A to the white B is deleted, B is marked. This prevents B from being GCed. The deletion barrier preserves weak tri-colour invariant: for any white object pointed to by black object, there exists some grey object from which through a series of white objects that white object is reachable. But when the pointer from the greyto the whiteis deleted,is marked. This preventsfrom being GCed. The deletion barrier preserves weak tri-colour invariant: for any white object pointed to by black object, there exists some grey object from which through a series of white objects that white object is reachable.

Deletion barrier is used by OCaml. We extend the write barrier to mark the original object in the reference r if both r and x are in the major heap. Deletion barrier is used by OCaml. We extend the write barrier to mark the original object in the referenceif bothandare in the major heap.

Let us extend the language by adding support for parallel execution. In Multicore OCaml, Domain.spawn forks off a new domain to run the thunk in parallel with the calling domain. It is reasonable to expect that most of the objects in the minor heap are in fact local to the domain which allocated the object, and are never shared between domains. Hence, it is desirable to collect each domain's young garbage independently without having to synchronize all of the domains. Let us extend the language by adding support for parallel execution. In Multicore OCaml,forks off a new domain to run the thunk in parallel with the calling domain. It is reasonable to expect that most of the objects in the minor heap are in fact local to the domain which allocated the object, and are never shared between domains. Hence, it is desirable to collect each domain's young garbage independently without having to synchronize all of the domains.

This can be done if the minor heap objects are only accessed by the owning domain. This can be done if the minor heap objects are only accessed by the owning domain.

In their In their POPL'93 paper , Doligez and Leroy built a concurrent garbage collector for concurrent caml light which used domain local heaps. In their paper, the heap invariants imposed are that there are no pointers between the local heaps and the major heap does not point to any minor heaps.

These heap invariants are enforced with the help of a write barrier that intercepts writes, and whenever there is a pointer about to be created between the major and a minor heap, the transitive closure of the minor heap object is promoted to the major heap before the assignment. These heap invariants are enforced with the help of a write barrier that intercepts writes, and whenever there is a pointer about to be created between the major and a minor heap, the transitive closure of the minor heap object is promoted to the major heap before the assignment.

The problem with this strategy is that we end up with false promotions. Consider a work stealing queue used for sharing work among multiple domains. It is very likely that the work-stealing queue survives the minor collection and is promoted to the major heap. Now, whenever a domain adds work to the queue, the work needs to be promoted to preserve the heap invariant, even though we expect that the domain which added the work to consume it in the common case. The problem with this strategy is that we end up with false promotions. Consider a work stealing queue used for sharing work among multiple domains. It is very likely that the work-stealing queue survives the minor collection and is promoted to the major heap. Now, whenever a domain adds work to the queue, the work needs to be promoted to preserve the heap invariant, even though we expect that the domain which added the work to consume it in the common case.

Hence, we relax this invariant to allow pointers from major heap to the minor. We still do not allow pointers between minor heaps, but allow pointers from major to minor heaps. Hence, we relax this invariant to allow pointers from major heap to the minor. We still do not allow pointers between minor heaps, but allow pointers from major to minor heaps.

We make sure that a domain does not access objects in a foreign domain with the help of a read barrier. The read barrier intercepts reads to mutable fields. If the value loaded is an integer, an object in the shared heap or own minor heap, then we let the read continue. Otherwise, we interrupt the domain which owns the object to promote the object closure. This operation returns the new location of the object in the major heap. This scheme ensures that the minor heaps can be independently GCed. We make sure that a domain does not access objects in a foreign domain with the help of a read barrier. The read barrier intercepts reads to mutable fields. If the value loaded is an integer, an object in the shared heap or own minor heap, then we let the read continue. Otherwise, we interrupt the domain which owns the object to promote the object closure. This operation returns the new location of the object in the major heap. This scheme ensures that the minor heaps can be independently GCed. Marlow and Peyton Jones evaluated a local heap design with similar weaker heap invariants for GHC Haskell.

There is a surprisingly efficient way to implement the read barrier checks through careful virtual memory mapping for minor heaps and a few bit twiddling tricks. Recall that the fast path for the read barrier is the value read is either an integer, or a shared heap value, or a value in own minor heap. Let us assume a 16-bit address space. The minor heap are all allocated in a contiguous power-of-2 aligned virtual memory area, where each minor heap is also a power-of-2 aligned and sized. Not all of the virtual memory area need to be allocated, but only needs to be reserved. In particular, we ensure that shared heap pages are not allocated in the minor heap area. There is a surprisingly efficient way to implement the read barrier checks through careful virtual memory mapping for minor heaps and a few bit twiddling tricks. Recall that the fast path for the read barrier is the value read is either an integer, or a shared heap value, or a value in own minor heap. Let us assume a 16-bit address space. The minor heap are all allocated in a contiguous power-of-2 aligned virtual memory area, where each minor heap is also a power-of-2 aligned and sized. Not all of the virtual memory area need to be allocated, but only needs to be reserved. In particular, we ensure that shared heap pages are not allocated in the minor heap area.

Addresses in this 16-bit address space can be written as 4 quads 0xPQRS . In OCaml, integer values are represented by tagging the least significant bit to be 1. Hence, in this example, integers have low bit of S to be 1. Minor heap values have PQ to be 42 , and R determines the domain. We can implement the read barrier check by comparing the given address with an address from the minor heap. Luckily, we have such an address handily available in the register -- the allocation pointer. On amd64, the allocation pointer is in register r15 . Addresses in this 16-bit address space can be written as 4 quads. In OCaml, integer values are represented by tagging the least significant bit to be 1. Hence, in this example, integers have low bit ofto be 1. Minor heap values haveto be, anddetermines the domain. We can implement the read barrier check by comparing the given address with an address from the minor heap. Luckily, we have such an address handily available in the register -- the allocation pointer. On amd64, the allocation pointer is in register

Here is our read barrier implementation for amd64 architecture. Assume that the value of interest is in rax register. At the end of this sequence of instructions, if none of the enabled bits in 0xff01 are set in rax , then zero flag will be set, and we know that the value is not a pointer into a foreign minor heap. Let's see how this works for the different cases. Here is our read barrier implementation for amd64 architecture. Assume that the value of interest is inregister. At the end of this sequence of instructions, if none of the enabled bits inare set in, then zero flag will be set, and we know that the value is not a pointer into a foreign minor heap. Let's see how this works for the different cases.

In the case, of integer, the least significant bit remains set, and hence zero flag will not be set. We are safe to read this value. In the case, of integer, the least significant bit remains set, and hence zero flag will not be set. We are safe to read this value.

In the case of a shared heap address, the PQ bits will different between r15 and rax . Hence, zero flag will not be set. In the case of a shared heap address, thebits will different betweenand. Hence, zero flag will not be set.

In the case of an address in own minor heap, the bits PQR will be the same. Hence, the subtraction underflows and sets all the bits in PQ . Hence, the zero flag will not be set. In the case of an address in own minor heap, the bitswill be the same. Hence, the subtraction underflows and sets all the bits in. Hence, the zero flag will not be set.

In the case of an address in foreign minor heap, the bits PQ will be the same. The bits in R will be different, and S will be zero in both values. After xoring, R will be non-zero and importantly, the rest of the bits are zero. Subtracting 1 from a non-zero value does not underflow, hence the rest of the bits remain zero. Now, the zero flag will be set after the test, and we know that the pointer is in the foreign minor heap. In the case of an address in foreign minor heap, the bitswill be the same. The bits inwill be different, andwill be zero in both values. After xoring,will be non-zero and importantly, the rest of the bits are zero. Subtracting 1 from a non-zero value does not underflow, hence the rest of the bits remain zero. Now, the zero flag will be set after the test, and we know that the pointer is in the foreign minor heap.

We now have an efficient way to find out whether we need to promote. But how do we perform the promotion on a read fault? The general strategy is to interrupt the foreign domain to perform the promotion for us. There are several alternatives here. On receiving the interrupt, the foreign domain may copy the objects from its minor heap to the major heap. While this strategy works for immutable objects, mutable objects which are copied should somehow be kept in sync on updates. This gets tricky especially with relaxed memory behaviours observed on modern multicore hardware. Moreover, this scheme breaks OCaml objects with Abstract_tag where a C library may map a C structure onto the OCaml heap and modify it transparently without the knowledge of the write barrier. Hence, the copies may go out of sync. We may instead move the object to the major heap and perform a minor GC to fix any references to the promoted objects. However, this scheme suffers from false promotions and long pauses during reads. To avoid false promotions, we may promote the object and scan the roots and the entire minor heap to fix any references to promoted objects. However, one needs to scan all the objects in the minor heap, which even the minor collection doesn't have to do.

We make the observation that most objects promoted on read faults happen to be recently allocated. The reason being that such objects are messages placed into channels where there is a waiting receiver on the other domain which can consume the message immediately. In our experiments on a small corpus of parallel OCaml programs, we found that 95% of objects promoted on read fault are among the youngest 5%. To make use of this fact, we combine the solutions 2 and 3 from above. We make the observation that most objects promoted on read faults happen to be recently allocated. The reason being that such objects are messages placed into channels where there is a waiting receiver on the other domain which can consume the message immediately. In our experiments on a small corpus of parallel OCaml programs, we found that 95% of objects promoted on read fault are among the youngest 5%. To make use of this fact, we combine the solutions 2 and 3 from above.

If the objected to be promoted on read fault is among the youngest x% (in the current implementation x = 10, but can be dynamically chosen), then we move the transitive object closure to the major heap. We may have extant pointers into the promoted objects. We need to fix those pointers such that they point to the new location of these objects in the major heap. In particular, the pointers may be found in the roots, objects younger than promoted object in the minor heap, and older minor objects that point to younger minor objects due to mutations of older objects. For the latter, we extend the write barrier to record such minor to minor pointers in promotion_set , which is only scanned during the promotion process. Otherwise, we move the object closure and perform a minor GC. If the objected to be promoted on read fault is among the youngest x% (in the current implementation x = 10, but can be dynamically chosen), then we move the transitive object closure to the major heap. We may have extant pointers into the promoted objects. We need to fix those pointers such that they point to the new location of these objects in the major heap. In particular, the pointers may be found in the roots, objects younger than promoted object in the minor heap, and older minor objects that point to younger minor objects due to mutations of older objects. For the latter, we extend the write barrier to record such minor to minor pointers in, which is only scanned during the promotion process. Otherwise, we move the object closure and perform a minor GC.

That resolves the major pain points in the interaction of parallelism with minor GC and promotion. The next is the interplay between parallelism and major GC. Recall that OCaml's major GC is incremental -- the mutator and the GC (mark and sweep) take turns to run. If we were to extend the scheme naively to parallel execution, we would have a stop-the-world incremental collector, where all the domains have to synchronize for GC work. This would introduce significant latency overheads. Instead, we go for a concurrent collector design where the mutator and the GC thread can run in parallel. That resolves the major pain points in the interaction of parallelism with minor GC and promotion. The next is the interplay between parallelism and major GC. Recall that OCaml's major GC is incremental -- the mutator and the GC (mark and sweep) take turns to run. If we were to extend the scheme naively to parallel execution, we would have a stop-the-world incremental collector, where all the domains have to synchronize for GC work. This would introduce significant latency overheads. Instead, we go for a concurrent collector design where the mutator and the GC thread can run in parallel.

Our design is based on the Our design is based on the Very Concurrent Mark-&-Sweep Garbage Collector (VCGC) design from the Inferno project which allows the mutator, marker and the sweeper threads to run concurrently. In VCGC design, there is a small stop-the-world phase at the end of a cycle where the threads agree on the end of the current major GC cycle and the beginning of the next one. Multicore OCaml's major GC is mostly concurrent mark and sweep where the stop-the-world phase might need to do a small fraction of major GC work left over before the end of the cycle, not unlike the VCGC design with many mutators i.e., parallel execution.

The major heap objects are in one of the 4 states: Marked , Unmarked , Garbage and Free . The domains alternate between running the mutator and performing GC work. The GC thread performs a depth-first traversal of the heap. If it finds an Unmarked object, it changes its colour to Marked and pushes the object into a domain-local mark-stack. On the other hand, if it finds a Garbage object, it marks it as Free and adds it to the free list. Since multiple GC threads operate on the heap simultaneously, marking is racy but idempotent. In particular, there is no synchronization for marking the objects. The major heap objects are in one of the 4 states:and. The domains alternate between running the mutator and performing GC work. The GC thread performs a depth-first traversal of the heap. If it finds anobject, it changes its colour toand pushes the object into a domain-local mark-stack. On the other hand, if it finds aobject, it marks it asand adds it to the free list. Since multiple GC threads operate on the heap simultaneously, marking is racy but idempotent. In particular, there is no synchronization for marking the objects.

If any domain thinks that all the work in the current major GC cycle is done (in practice, close to being done), it calls for a global synchronization where all the domains synchronize on the barrier. Once stopped, all the domains work to actually finish marking, as some work may be left over at other domains. Once that is done, the cores agree to flip the meaning of the colours. Anything that is Unmarked is considered Garbage . Anything that is Marked becomes Unmarked for the next cycle. Garbage objects are considered Marked , but at the end of the major GC, all Garbage objects have been marked Free . Hence, no objects fall into this category. Anything that is marked Free still remains Free . This concludes the discussion on parallelism. If any domain thinks that all the work in the current major GC cycle is done (in practice, close to being done), it calls for a global synchronization where all the domains synchronize on the barrier. Once stopped, all the domains work to actually finish marking, as some work may be left over at other domains. Once that is done, the cores agree to flip the meaning of the colours. Anything that isis considered. Anything that isbecomesfor the next cycle.objects are considered, but at the end of the major GC, allobjects have been marked. Hence, no objects fall into this category. Anything that is markedstill remains. This concludes the discussion on parallelism.

We introduce concurrency into the mix next. Concurrency in Multicore OCaml is expressed through fibers, which are language-level lightweight threads. Fibers are implemented as heap allocated, dynamically resized stack segments. Just like mutating a regular heap object, fibers can also be mutated by pushing and popping values. However, unlike regular objects, fiber mutations are not protected by a write barrier. This poses challenges to maintaining the heap invariants. We introduce concurrency into the mix next. Concurrency in Multicore OCaml is expressed through fibers, which are language-level lightweight threads. Fibers are implemented as heap allocated, dynamically resized stack segments. Just like mutating a regular heap object, fibers can also be mutated by pushing and popping values. However, unlike regular objects, fiber mutations are not protected by a write barrier. This poses challenges to maintaining the heap invariants.

The main thread in Multicore OCaml is also a fiber. Hence, the GC root current stack is just a pointer to the current fiber. Since we don't have write barrier on pushing to a fiber, we need to approximate the pointers that may arise from a fiber in the major heap which points to the minor heap. We do this with the help of remembered fiber set. The main thread in Multicore OCaml is also a fiber. Hence, the GC root current stack is just a pointer to the current fiber. Since we don't have write barrier on pushing to a fiber, we need to approximate the pointers that may arise from a fiber in the major heap which points to the minor heap. We do this with the help of remembered fiber set.

This represents the set of fibers in the major heap that ran in the current minor cycle on some domain. Similar to remembered set, the remembered fiber set is also a domain local data structure. The remembered fiber set is a root for minor GCs. The remembered fiber set is cleared along with the remembered set at the end of minor GC. This represents the set of fibers in the major heap that ran in the current minor cycle on some domain. Similar to remembered set, the remembered fiber set is also a domain local data structure. The remembered fiber set is a root for minor GCs. The remembered fiber set is cleared along with the remembered set at the end of minor GC.

We also treat fibers specially during promotions on read faults. Fibers transitively reachable are not promoted automatically in order to avoid prematurely promoting the entire world. We envision work stealing schedulers where the fiber only needs to be promoted when a different domain explicitly demands it. In this example, when assigning x to r , instead of promoting fiber f , In this example, when assigningto, instead of promoting fiber

we leave it on the minor heap. We record x in the remembered set. we leave it on the minor heap. We recordin the remembered set.

And when a different domain demands it i.e., tries to context switch to the fiber f , And when a different domain demands it i.e., tries to context switch to the fiber

then we promote the fiber and its transitive closure. then we promote the fiber and its transitive closure.

Since we've added the domain-local remembered fiber set, at every instance the fast path is taken on a read fault, we are obligated to scan this set just after the promotion is done. But the remembered fiber set may be large, and the fiber stacks themselves can be large. Hence, scanning the stack for most promotions would introduce large pause times. Since we've added the domain-local remembered fiber set, at every instance the fast path is taken on a read fault, we are obligated to scan this set just after the promotion is done. But the remembered fiber set may be large, and the fiber stacks themselves can be large. Hence, scanning the stack for most promotions would introduce large pause times.

We would like to break up the large pause time if possible. We can do this by lazily scanning the fibers just before a context switch. We only need to scan a fiber once per promotion. We also expect the rate of context switches to be much smaller than the rate of promotions. Hence, in practice, the fiber only gets scanned once per batch of promotions. We would like to break up the large pause time if possible. We can do this by lazily scanning the fibers just before a context switch. We only need to scan a fiber once per promotion. We also expect the rate of context switches to be much smaller than the rate of promotions. Hence, in practice, the fiber only gets scanned once per batch of promotions.

How does concurrency affect the major GC? Recall that Multicore OCaml uses a deletion barrier. In a deletion barrier, whenever a major heap object loses a reference to another major heap object, we have to mark the target object. However, fiber stack pops are not protected by a write barrier. Instead, we conservatively mark all of the objects on the fiber before switching to it. This is not to dissimilar to the current stock OCaml compiler scanning the current stack (as part of the roots) at the beginning of the major GC cycle. How does concurrency affect the major GC? Recall that Multicore OCaml uses a deletion barrier. In a deletion barrier, whenever a major heap object loses a reference to another major heap object, we have to mark the target object. However, fiber stack pops are not protected by a write barrier. Instead, we conservatively mark all of the objects on the fiber before switching to it. This is not to dissimilar to the current stock OCaml compiler scanning the current stack (as part of the roots) at the beginning of the major GC cycle.

In VCGC, marking is racy but idempotent. Hence, the mutators can freely read and write the object while it is concurrently marked. In particular, no synchronization such as compare-and-swap or locks are required to mediate access between the mutator and the GC thread. However, this invariant does not hold for fibers. Assume that a mutator is about to context switch to a thread while a GC thread on another domain attempts to scan and mark the objects on the fiber. Since the mutator may push and pop the fiber stack, the GC thread concurrently scanning the stack may observe inconsistent state. In VCGC, marking is racy but idempotent. Hence, the mutators can freely read and write the object while it is concurrently marked. In particular, no synchronization such as compare-and-swap or locks are required to mediate access between the mutator and the GC thread. However, this invariant does not hold for fibers. Assume that a mutator is about to context switch to a thread while a GC thread on another domain attempts to scan and mark the objects on the fiber. Since the mutator may push and pop the fiber stack, the GC thread concurrently scanning the stack may observe inconsistent state.

In order to prevent this inconsistency, before switching to a unmarked fiber, the fiber is marked and all the objects on the fiber are also marked -- fiber is made Black . And in order to prevent racy access, the fiber is locked while marking. If the GC thread holds the lock on a fiber and a mutator tries to context switch to it, the mutator blocks until the fiber is marked. If the GC thread loses the race, it can safely skip marking the fiber. In order to prevent this inconsistency, before switching to a unmarked fiber, the fiber is marked and all the objects on the fiber are also marked -- fiber is made. And in order to prevent racy access, the fiber is locked while marking. If the GC thread holds the lock on a fiber and a mutator tries to context switch to it, the mutator blocks until the fiber is marked. If the GC thread loses the race, it can safely skip marking the fiber.

In summary, the Multicore OCaml GC optimizes for latency and minimizes the maximum pause time through independent minor GCs and a mostly concurrent mark-and-sweep collector for major GCs. The challenge is to consider the interactions between mutations, concurrency and parallelism and the various GC techniques in order to come up with a design that preserves safety and optimizes our performance goals. I will have to save the performance analysis of the GC to a different talk. In summary, the Multicore OCaml GC optimizes for latency and minimizes the maximum pause time through independent minor GCs and a mostly concurrent mark-and-sweep collector for major GCs. The challenge is to consider the interactions between mutations, concurrency and parallelism and the various GC techniques in order to come up with a design that preserves safety and optimizes our performance goals. I will have to save the performance analysis of the GC to a different talk.

Further Reading