While on vacation, I’ve been working on an alternate type inference scheme for rustc. (Actually, I got it 99% working on the plane, and have been slowly poking at it ever since.) This scheme simplifies the code of the type inferencer dramatically and (I think) helps to meet our intutions (as I will explain). It is however somewhat less flexible than the existing inference scheme, though all of rustc and all the libraries compile without any changes. The scheme will (I believe) make it much simpler to implement to proper one-way matching for traits (explained later).

Note: Changing the type inference scheme doesn’t really mean much to end users. Roughly the same set of Rust code still compiles. So this post is really mostly of interest to rustc implementors.

The new scheme in a nutshell

The new scheme is fairly simple. It is based on the observation that most subtyping in Rust arises from lifetimes (though the scheme is extensible to other possible kinds of subtyping, e.g. virtual structs). It abandons unification and the H-M infrastructure and takes a different approach: when a type variable V is first related to some type T , we don’t set the value of V to T directly. Instead, we say that V is equal to some type U where U is derived by replacing all lifetimes in T with lifetime variables. We then relate T and U appropriately.

Let me give an example. Here are two variables whose type must be inferred:

'a: { // 'a --> name of block's lifetime let x = 3; let y = &x; ... }

Let’s say that the type of x is $X and the type of y is $Y , where $X and $Y are both inference variables. In that case, the first assignment generates the constraint that int <: $X and the second generates the constraint that &'a $X <: $Y . To resolve the first constraint, we would set $X directly to int . This is because there are no lifetimes in the type int . To resolve the second constraint, we would set $Y to &'0 int – here '0 represents a fresh lifetime variable. We would then say that &'a int <: &'0 int , which in turn implies that '0 <= 'a . After lifetime inference is complete, the types of x and y would be int and &'a int as expected.

Without unification, you might wonder what happens when two type variables are related that have not yet been associated with any concrete type. This is actually somewhat challenging to engineer, but it certainly does happen. For example, there might be some code like:

let mut x; // type: $X let mut y = None; // type: Option<$0> loop { if y.is_some() { x = y.unwrap(); ... } ... }

Here, at the point where we process x = y.unwrap() , we do not yet know the values of either $X or $0 . We can say that the type of y.unwrap() will be $0 but we must now process the constrint that $0 <: $X . We do this by simply keeping a list of outstanding constraints. So neither $0 nor $X would (yet) be assigned a specific type, but we’d remember that they were related. Then, later, when either $0 or $X is set to some specific type T , we can go ahead and instantiate the other with U , where U is again derived from T by replacing all lifetimes with lifetime variables. Then we can relate T and U appropriately.

If we wanted to extend the scheme to handle more kinds of inference beyond lifetimes, it can be done by adding new kinds of inference variables. For example, if we wanted to support subtyping between structs, we might add struct variables.

What advantages does this scheme have to offer?

The primary advantage of this scheme is that it is easier to think about for us compiler engineers. Every type variable is either set – in which case its type is known precisely – or unset – in which case its type is not known at all. In the current scheme, we track a lower- and upper-bound over time. This makes it hard to know just how much is really known about a type. Certainly I know that when I think about inference I still think of the state of a variable as a binary thing, even though I know that really it’s something which evolves.

What prompted me to consider this redesign was the need to support one-way matching as part of trait resolution. One-way matching is basically a way of saying: is there any substitution S such that T <: S(U) (whereas normal matching searches for a substitution applied to both sides, like S(T) <: S(U) ).

One-way matching is very complicated to support in the current inference scheme: after all, if there are type variables that appear in T or U which are partially constrained, we only know bounds on their eventual type. In practice, these bounds actually tell us a lot: for example, if a type variable has a lower bound of int , it actually tells us that the type variable is int , since in Rust’s type system there are no super- of sub-types of int . However, encoding this sort of knowledge is rather complex – and ultimately amounts to precisely the same thing as this new inference scheme.

Another advantage is that there are various places in the Rust’s type checker whether we query the current state of a type variable and make decisions as a result. For example, when processing *x , if the type of x is a type variable T , we would want to know the current state of T – is T known to be something inherent derefable (like &U or &mut U ) or a struct that must implement the Deref trait? The current APIs for doing this bother me because they expose the bounds of U – but those bounds can change over time. This seems “risky” to me, since it’s only sound for us to examine those bounds if we either (a) freeze the type of T or (b) are certain that we examine properties of the bound that will not change. This problem does not exist in the new inference scheme: anything that might change over time is abstracted into a new inference variable of its own.

What are the disadvantages?

One form of subtyping that exists in Rust is not amenable to this inference. It has to do with universal quantification and function types. Function types that are “more polymorphic” can be subtypes of functions that are “less polymorphic”. For example, if I have a function type like <'a> fn(&'a T) -> &'a uint , this indicates a function that takes a reference to T with any lifetime 'a and returns a reference to a uint with that same lifetime. This is a subtype of the function type fn(&'b T) -> &'b uint . While these two function types look similar, they are quite different: the former accepts a reference with any lifetime but the latter accepts only a reference with the specific lifetime 'b .

What this means is that today if you have a variable that is assigned many times from functions with varying amounts of polymorphism, we will generally infer its type correctly:

fn example<'b>(..) { let foo: <'a> |&'a T| -> &'a int = ...; let bar: |&'b T| -> &'b int = ...; let mut v; v = foo; v = bar; // type of v is inferred to be |&'b T| -> &'b int }

However, this will not work in the newer scheme. Type ascription of some form would be required. As you can imagine, this is not a very .common problem, and it did not arise in any existing code.

(I believe that there are situations which the newer scheme infers correct types and the older scheme will fail to compile; however, I was unable to come up with a good example.)

How does it perform?

I haven’t done extensive measurements. The newer scheme creates a lot of region variables. It seems to perform roughly the same as the older scheme, perhaps a bit slower – optimizing region inference may be able to help.