At the recent compiler design sprint, we spent some time discussing non-lexical lifetimes, the plan to make Rust’s lifetime system significantly more advanced. I want to write-up those plans here, and give some examples of the kinds of programs that would now type-check, along with some that still will not (for better or worse).

If you were at the sprint, then the system I am going to describe in this blog post will actually sound quite a bit different than what we were talking about. However, I believe it is equivalent to that system. I am choosing to describe it differently because this version, I believe, would be significantly more efficient to implement (if implemented naively). I also find it rather easier to understand.

I have a prototype implementation of this system. The example used in this post, along with the ones from previous posts, have all been tested in this prototype and work as expected.

Yet another example

I’ll start by giving an example that illustrates the system pretty well, I think. This section also aims to give an intution for how the system works and what set of programs will be accepted without going into any of the details. Somewhat oddly, I’m going to number this example as “Example 4”. This is because my previous post introduced examples 1, 2, and 3. If you’ve not read that post, you may want to, but don’t feel you have to. The presentation in this post is intended to be independent.

Example 4: Redefined variables and liveness

I think the key ingredient to understanding how NLL should work is understanding liveness. The term “liveness” derives from compiler analysis, but it’s fairly intuitive. We say that a variable is live if the current value that it holds may be used later. This is very important to Example 4:

let mut foo , bar ; let p = & foo ; // `p` is live here: its value may be used on the next line. if condition { // `p` is live here: its value will be used on the next line. print ( * p ); // `p` is DEAD here: its value will not be used. p = & bar ; // `p` is live here: its value will be used later. } // `p` is live here: its value may be used on the next line. print ( * p ); // `p` is DEAD here: its value will not be used.

Here you see a variable p that is assigned in the beginning of the program, and then maybe re-assigned during the if . The key point is that p becomes dead (not live) in the span before it is reassigned. This is true even though the variable p will be used again, because the value that is in p will not be used.

So how does liveness relate to non-lexical lifetimes? The key rule is this: Whenever a variable is live, all references that it may contain are live. This is actually a finer-grained notion than just the liveness of a variable, as we will see. For example, the first assignment to p is &foo – we want foo to be borrowed everywhere that this assignment may later be accessed. This includes both print() calls, but excludes the period after p = &bar . Even though the variable p is live there, it now holds a different reference:

let foo , bar ; let p = & foo ; // `foo` is borrowed here, but `bar` is not if condition { print ( * p ); // neither `foo` nor `bar` are borrowed here p = & bar ; // assignment 1 // `foo` is not borrowed here, but `bar` is } // both `foo` and `bar` are borrowed here print ( * p ); // neither `foo` nor `bar` are borrowed here, // as `p` is dead

Our analysis will begin with the liveness of a variable (the coarser-grained notion I introduced first). However, it will use reachability to refine that notion of liveness to obtain the liveness of individual values.

Control-flow graphs and point notation

Recall that in NLL-land, all reasoning about lifetimes and borrowing will take place in the context of MIR, in which programs are represented as a control-flow graph. This is what Example 4 looks like as a control-flow graph:

// let mut foo: i32; // let mut bar: i32; // let p: &i32; A [ p = &foo ] [ if condition ] ----\ (true) | | | B v | [ print(*p) ] | [ ... ] | [ p = &bar ] | [ ... ] | [ goto C ] | | +-------------/ | C v [ print(*p) ] [ return ]

As a reminder, I will use a notation like Block/Index to refer to a specific point (statement) in the control-flow graph. So A/0 and B/2 refer to p = &foo and p = &bar , respectively. Note that there is also a point for the goto/return terminators of each block (i.e., A/1, B/4, and C/1).

Using this notation, we can say that we want foo to be borrowed during the points A/1, B/0, and C/0. We want bar to be borrowed during the points B/3, B/4, and C/0.

Defining the NLL analysis

Now that we have our two examples, let’s work on defining how the NLL analysis will work.

Step 0: What is a lifetime?

The lifetime of a reference is defined in our system to be a region of the control-flow graph. We will represent such regions as a set of points.

A note on terminology: For the remainder of this post, I will often use the term region in place of “lifetime”. Mostly this is because it’s the standard academic term and it’s often the one I fall back to when thinking more formally about the system, but it also feels like a good way to differentiate the lifetime of the reference (the region where it is in use) with the lifetime of the referent (the span of time before the underlying resource is freed).

Step 1: Instantiate erased regions

The plan for adopting NLL is to do type-checking in two phases. The first phase, which is performed on the HIR, I would call type elaboration. This is basically the “traditional type-system” phase. It infers the types of all variables and other things, figures out where autoref goes, and so forth; the result of this is the MIR.

The key change from today is that I want to do all of this type elaboration using erased regions. That is, until we build the MIR, we won’t have any regions at all. We’ll just keep a placeholder (which I’ll write as 'erased ). So if you have something like &i32 , the elaborated, internal form would just be &'erased i32 . This is quite different from today, where the elaborated form includes a specific region. (However, this erased form is precisely what we want for generating code, and indeed MIR today goes through a “region erasure” step; this step would be unnecessary in the new plan, since MIR as produced by type check would always have fully erased regions.)

Once we have built MIR, then, the idea is roughly to go and replace all of these erased regions with inference variables. This means we’ll have region inference variables in the types of all local variables; it also means that for each borrow expression like &foo , we’ll have a region representing the lifetime of the resulting reference. I’ll write the expression together with this region like so: &'0 foo .

Here is what the CFG for Example 4 looks like with regions instantiated. You can see I used the variable '0 to represent the region in the type of p , and '1 and '2 for the regions of the two borrows:

// let mut foo: i32; // let mut bar: i32; // let p: &'0 i32; A [ p = &'1 foo ] [ if condition ] ----\ (true) | | | B v | [ print(*p) ] | [ ... ] | [ p = &'2 bar ] | [ ... ] | [ goto C ] | | +-------------/ | C v [ print(*p) ] [ return ]

Step 2: Introduce region constraints

Now that we have our region variables, we have to introduce constraints. These constriants will come in two kinds:

liveness constraints; and,

subtyping constraints.

Let’s look at each in turn.

Liveness constraints.

The basic rule is this: if a variable is live on entry to a point P, then all regions in its type must include P.

Let’s continue with Example 4. There, we have just one variable, p . It’s type has one region ( '0 ) and it is live on entry to A/1, B/0, B/3, B/4, and C/0. So we wind up with a constraint like this:

{A/1, B/0, B/3, B/4, C/0} <= '0

We also include a rule that for each borrow expression like &'1 foo , '1 must include the point of borrow. This gives rise to two further constraints in Example 4:

{A/0} <= '1 {B/2} <= '2

Location-aware subtyping constraints

The next thing we do is to go through the MIR and establish the normal subtyping constraints. However, we are going to do this with a slight twist, which is that we are going to take the current location into account. That is, instead of writing T1 <: T2 ( T1 is required to be a subtype of T2 ) we will write (T1 <: T2) @ P ( T1 is required to be a subtype of T2 at the point P). This in turn will translate to region constraints like (R2 <= R1) @ P .

Continuing with Example 4, there are a number of places where subtyping constraints arise. For example, at point A/0, we have p = &'1 foo . Here, the type of &'1 foo is &'1 i32 , and the type of p is &'0 i32 , so we have a (location-aware) subtyping constraint:

(&'1 i32 <: &'0 i32) @ A/1

which in turn implies

('0 <= '1) @ A/1 // Note the order is reversed.

Note that the point here is A/1, not A/0. This is because A/1 is the first point in the CFG where this constraint must hold on entry.

The meaning of a region constraint like ('0 <= '1) @ P is that, starting from the point P, the region '1 must include all points that are reachable without leaving the region '0 . The implementation basically does a depth-first search starting from P; the search stops if we exit the region '0 . Otherwise, for each point we find, we add it to '1 .

Jumping back to example 4, we wind up with two constraints in total. Combining those with the liveness constraint, we get this:

('0 <= '1) @ A/1 ('0 <= '2) @ B/3 {A/1, B/0, B/3, B/4, C/0} <= '0 {A/0} <= '1 {B/2} <= '2

We can now try to find the smallest values for '0 , '1 , and '2 that will make this true. The result is:

'0 = {A/1, B/0, B/3, B/4, C/0} '1 = {A/0, A/1, B/0, C/0} '2 = {B/3, B/4, C/0}

These results are exactly what we wanted. The variable foo is borrowed for the region '1 , which does not include B/3 and B/4. This is true even though the '0 includes those points; this is because you cannot reach B/3 and B/4 from A/1 without going through B/1, and '0 does not include B/1 (because p is not live at B/1). Similarly, bar is borrowed for the region '2 , which begins at B/4 and extends to C/0 (and need not include earlier points, which are not reachable).

You may wonder why we do not have to include all points in '0 in '1 . Intuitively, the reasoning here is based on liveness: '1 must ultimately include all points where the reference may be accessed. In this case, the subregion constraint arises because we are copying a reference (with region '1 ) into a variable (let’s call it x ) whose type includes the region '0 , so we need reads of '0 to also be counted as reads of '1 – but, crucially, only those reads that may observe this write. Because of the liveness constraints we saw earlier, if x will later be read, then x must be live along the path from this copy to that read (by the definition of liveness, essentially). Therefore, because the variable is live, '0 will include that entire path. Hence, by including the points in '0 that are reachable from the copy (without leaving '0 ), we include all potential reads of interest.

Conclusion

This post presents a system for computing non-lexical lifetimes. It assumes that all regions are erased when MIR is created. It uses only simple compiler concepts, notably liveness, but extends the subtyping relation to take into account where the subtyping must hold. This allows it to disregard unreachable portions of the control-flow.

I feel pretty good about this iteration. Among other things, it seems so simple I can’t believe it took me this long to come up with it. This either means that is it the right thing or I am making some grave error. If it’s the latter people will hopefully point it out to me. =) It also seems to be efficiently implementable.

I want to emphasize that this system is the result of a lot of iteration with a lot people, including (but not limited to) Cameron Zwarich, Ariel Ben-Yehuda, Felix Klock, Ralf Jung, and James Miller.

It’s interesting to compare this with various earlier attempts:

Our earliest thoughts assumed continuous regions (e.g., RFC 396). The idea was that the region for a reference ought to correspond to some continuous bit of control-flow, rather than having “holes” in the middle. The example in this post shows the limitation of this, however. Note that the region for the variable p includes B/0 and B/4 but excludes B/1. This is why we lean on liveness requirements instead, so as to ensure that the region contains all paths from where a reference is created to where it is eventually dereferenced.

An alternative solution might be to consider continuous regions but apply an SSA or SSI transform. This allows the example in this post to type, but it falls down on more advanced examples, such as vec-push-ref (hat tip, Cameron Zwarich). In particular, it’s possible for subregion relations to arise without a variable being redefined. You can go farther, and give variables a distinct type at each point in the program, as in Ericson2314’s stateful MIR for Rust. But even then you must contend with invariance or you have the same sort of problems. Exploring this led to the development of the “localized” subregion relationship constraint (r1 <= r2) @ P , which I had in mind in my original series but which we elaborated more fully at the rustc design sprint. The change in this post versus what we said at the sprint is that I am using one type per variable instead of one type per variable per statement; I am also explicitly using the results of an earlier liveness analysis to construct the constraints, whereas in the sprint we incorporated the liveness into the region inference itself (by reasoning about which values were live across each individual statement and thus creating many more inequalities).



There are some things I’ve left out of this post. Hopefully I will get to them in future posts, but they all seem like relatively minor twists on this base model.

I’d like to talk about how to incorporate lifetime parameters on fns (I think we can do that in a fairly simple way by modeling them as regions in an expanded control-flow graph, as illustrated by this example in my prototype).

There are various options for modeling the “deferred borrows” needed to accept vec.push(vec.len()) .

. We might consider a finer-grained notion of liveness that operates not on variables but rather on the “fragments” (paths) that we use when doing move-checking. This would help to make let (p, q) = (&foo, &bar) and let pair = (&foo, &bar) entirely equivalent (in the system as I described it, they are not, because whenever pair is live, both foo and bar would be borrowed, even if only pair.0 is ever used). But even if we do this there will still be cases where storing pointers into an aggregate (e.g., a struct) can lose precision versus using variables on the stack, so I’m not sure it’s worth worrying about.

Comments? Let’s use this old internals thread.