WARNING: This README is obsolete and will be removed soon! For more info on how the current borrowck works, see the rustc guide.
As of edition 2018, region inference is done using Non-lexical lifetimes, which is described in the guide and this RFC.
Note that we use the terms region and lifetime interchangeably.
Region inference uses a somewhat more involved algorithm than type inference. It is not the most efficient thing ever written though it seems to work well enough in practice (famous last words). The reason that we use a different algorithm is because, unlike with types, it is impractical to hand-annotate with regions (in some cases, there aren't even the requisite syntactic forms). So we have to get it right, and it's worth spending more time on a more involved analysis. Moreover, regions are a simpler case than types: they don't have aggregate structure, for example.
Basically our input is a directed graph where nodes can be divided
into two categories: region variables and concrete regions. Each edge
R -> S in the graph represents a constraint that the region R is a
subregion of the region S.
Region variable nodes can have arbitrary degree. There is one region variable node per region variable.
Each concrete region node is associated with some, well, concrete region: e.g., a free lifetime, or the region for a particular scope. Note that there may be more than one concrete region node for a particular region value. Moreover, because of how the graph is built, we know that all concrete region nodes have either in-degree 1 or out-degree 1.
Before resolution begins, we build up the constraints in a hashmap
that maps Constraint keys to spans. During resolution, we construct
the actual Graph structure that we describe here.
The algorithm is a simple dataflow algorithm. Each region variable
begins as empty. We iterate over the constraints, and for each constraint
we grow the relevant region variable to be as big as it must be to meet all the
constraints. This means the region variables can grow to be 'static if
necessary.
After all constraints are fully propoagated, we do a "verification" step where we walk over the verify bounds and check that they are satisfied. These bounds represent the "maximal" values that a region variable can take on, basically.
Let's first consider the region hierarchy without thinking about closures, because they add a lot of complications. The region hierarchy basically mirrors the lexical structure of the code. There is a region for every piece of 'evaluation' that occurs, meaning every expression, block, and pattern (patterns are considered to "execute" by testing the value they are applied to and creating any relevant bindings). So, for example:
fn foo(x: isize, y: isize) { // -+
// +------------+ // |
// | +-----+ // |
// | +-+ +-+ +-+ // |
// | | | | | | | // |
// v v v v v v v // |
let z = x + y; // |
... // |
} // -+
fn bar() { ... }In this example, there is a region for the fn body block as a whole,
and then a subregion for the declaration of the local variable.
Within that, there are sublifetimes for the assignment pattern and
also the expression x + y. The expression itself has sublifetimes
for evaluating x and y.
#s## Function calls
Function calls are a bit tricky. I will describe how we handle them now and then a bit about how we can improve them (Issue #6268).
Consider a function call like func(expr1, expr2), where func,
arg1, and arg2 are all arbitrary expressions. Currently,
we construct a region hierarchy like:
+----------------+
| |
+--+ +---+ +---+|
v v v v v vv
func(expr1, expr2)
Here you can see that the call as a whole has a region and the
function plus arguments are subregions of that. As a side-effect of
this, we get a lot of spurious errors around nested calls, in
particular when combined with &mut functions. For example, a call
like this one
self.foo(self.bar())where both foo and bar are &mut self functions will always yield
an error.
Here is a more involved example (which is safe) so we can see what's going on:
struct Foo { f: usize, g: usize }
// ...
fn add(p: &mut usize, v: usize) {
*p += v;
}
// ...
fn inc(p: &mut usize) -> usize {
*p += 1; *p
}
fn weird() {
let mut x: Box<Foo> = box Foo { /* ... */ };
'a: add(&mut (*x).f,
'b: inc(&mut (*x).f)) // (..)
}The important part is the line marked (..) which contains a call to
add(). The first argument is a mutable borrow of the field f. The
second argument also borrows the field f. Now, in the current borrow
checker, the first borrow is given the lifetime of the call to
add(), 'a. The second borrow is given the lifetime of 'b of the
call to inc(). Because 'b is considered to be a sublifetime of
'a, an error is reported since there are two co-existing mutable
borrows of the same data.
However, if we were to examine the lifetimes a bit more carefully, we
can see that this error is unnecessary. Let's examine the lifetimes
involved with 'a in detail. We'll break apart all the steps involved
in a call expression:
'a: {
'a_arg1: let a_temp1: ... = add;
'a_arg2: let a_temp2: &'a mut usize = &'a mut (*x).f;
'a_arg3: let a_temp3: usize = {
let b_temp1: ... = inc;
let b_temp2: &'b = &'b mut (*x).f;
'b_call: b_temp1(b_temp2)
};
'a_call: a_temp1(a_temp2, a_temp3) // (**)
}Here we see that the lifetime 'a includes a number of substatements.
In particular, there is this lifetime I've called 'a_call that
corresponds to the actual execution of the function add(), after
all arguments have been evaluated. There is a corresponding lifetime
'b_call for the execution of inc(). If we wanted to be precise
about it, the lifetime of the two borrows should be 'a_call and
'b_call respectively, since the references that were created
will not be dereferenced except during the execution itself.
However, this model by itself is not sound. The reason is that while the two references that are created will never be used simultaneously, it is still true that the first reference is created before the second argument is evaluated, and so even though it will not be dereferenced during the evaluation of the second argument, it can still be invalidated by that evaluation. Consider this similar but unsound example:
struct Foo { f: usize, g: usize }
// ...
fn add(p: &mut usize, v: usize) {
*p += v;
}
// ...
fn consume(x: Box<Foo>) -> usize {
x.f + x.g
}
fn weird() {
let mut x: Box<Foo> = box Foo { ... };
'a: add(&mut (*x).f, consume(x)) // (..)
}In this case, the second argument to add actually consumes x, thus
invalidating the first argument.
So, for now, we exclude the call lifetimes from our model.
Eventually I would like to include them, but we will have to make the
borrow checker handle this situation correctly. In particular, if
there is a reference created whose lifetime does not enclose
the borrow expression, we must issue sufficient restrictions to ensure
that the pointee remains valid.
Integrating closures properly into the model is a bit of work-in-progress. In an ideal world, we would model closures as closely as possible after their desugared equivalents. That is, a closure type would be modeled as a struct, and the region hierarchy of different closure bodies would be completely distinct from all other fns. We are generally moving in that direction but there are complications in terms of the implementation.
In practice what we currently do is somewhat different. The basis for the current approach is the observation that the only time that regions from distinct fn bodies interact with one another is through an upvar or the type of a fn parameter (since closures live in the fn body namespace, they can in fact have fn parameters whose types include regions from the surrounding fn body). For these cases, there are separate mechanisms which ensure that the regions that appear in upvars/parameters outlive the dynamic extent of each call to the closure:
- Types must outlive the region of any expression where they are used.
For a closure type
Cto outlive a region'r, that implies that the types of all its upvars must outlive'r. - Parameters must outlive the region of any fn that they are passed to.
Therefore, we can -- sort of -- assume that any region from an
enclosing fns is larger than any region from one of its enclosed
fn. And that is precisely what we do: when building the region
hierarchy, each region lives in its own distinct subtree, but if we
are asked to compute the LUB(r1, r2) of two regions, and those
regions are in disjoint subtrees, we compare the lexical nesting of
the two regions.
Ideas for improving the situation: (FIXME #3696) The correctness argument here is subtle and a bit hand-wavy. The ideal, as stated earlier, would be to model things in such a way that it corresponds more closely to the desugared code. The best approach for doing this is a bit unclear: it may in fact be possible to actually desugar before we start, but I don't think so. The main option that I've been thinking through is imposing a "view shift" as we enter the fn body, so that regions appearing in the types of fn parameters and upvars are translated from being regions in the outer fn into free region parameters, just as they would be if we applied the desugaring. The challenge here is that type inference may not have fully run, so the types may not be fully known: we could probably do this translation lazilly, as type variables are instantiated. We would also have to apply a kind of inverse translation to the return value. This would be a good idea anyway, as right now it is possible for free regions instantiated within the closure to leak into the parent: this currently leads to type errors, since those regions cannot outlive any expressions within the parent hierarchy. Much like the current handling of closures, there are no known cases where this leads to a type-checking accepting incorrect code (though it sometimes rejects what might be considered correct code; see #22557), but it still doesn't feel like the right approach.