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The compiler

Inko's compiler compiles Inko source files into a bytecode image, run by Inko's VM. The compiler is written in Rust. Originally we had plans for a self-hosting compiler, but have decided against taking this path (for the time being) to simplify the process of developing and maintaining Inko.

The compilation process consists of the following steps:

  1. Source code is recursively parsed into an AST, following any imported modules.
  2. The AST is lowered into Inko's high-level intermediate representation, known as "HIR".
  3. Type-checking is performed using HIR.
  4. HIR is lowered into Inko's mid-level intermediate representation, creatively known as "MIR". MIR is a graph based linear IR, instead of the tree based HIR/AST.
  5. MIR is used for enforcing single ownership and move semantics, optimisations, and other analysis that requires a more detailed IR (compared to HIR).
  6. MIR is lowered into Inko bytecode, which is then written to a bytecode image.

Parsing and the AST

Parsing is a topic extensively covered in computer science, so we won't go over it in much detail. Inko uses a handwritten LL(1) recursive descend parser, and a separate lexer/tokeniser. Each type of thing to parse (e.g. a class or an array literal) typically has a corresponding function in the parser. AST nodes are just plain structures.

Each AST node stores a source location, which is a line and column range. Source file paths are not stored per AST node, instead we implicitly assume that all AST nodes originating from a module share the same file; removing the need for duplicating the file path for every AST node.


HIR is essentially just the AST with some minor changes, and the ability to store type information in nodes. Type information is typically stored in an optional field called resolved_type (defaulting to None), though some nodes may use different fields.

HIR exists because we want to keep the AST as close as possible to the source code, making it possible to use the AST for e.g. a code formatter, but this leads to it being a bit annoying to work with when type checking.

Type checking

The type checker is essentially a recursive descend checker. It traverses the HIR nodes and type-checks them in a depth-first order. Nodes are annotated with their typing information whenever necessary.

Typing information is stored in the "type database". This is a structure that looks a bit like a relational database, and is addressed using indexes to arrays instead of pointers/references. It's these IDs that are stored in the AST nodes. These IDs are globally unique, including for local variables. This makes it easier for MIR to inline code, as MIR doesn't care about the variable names and instead uses the variable IDs.

Type checking is done by performing multiple passes over the HIR, though not all passes transform the entire HIR. Passes are typically just structures with a run_all method that operates on a list of modules to check. For example, there's a type checking pass that just defines types without descending into their methods, while another pass ensures type parameter requirements are valid. The result is a setup somewhat inspired by The Nanopass Framework, though not as extreme.

In general the compiler follows the principal that it's better to have too many passes than too little, as merging passes is easier compared to splitting them up.


MIR is used for enforcing single ownership, move semantics, and more. It's a graph based IR instead of a tree based IR. The graph consists of structures known as basic blocks. Each block has zero or more instructions, stored in an array. A block may flow into other blocks known as successors. The blocks that flow into other blocks are known as the predecessor, as in A -> B means A is the predecessor of B, while B is the successor of A. A block always ends in an explicit jump to the next block, allowing the compiler to process blocks in any order, making code generation easier.

Pattern matching

Match expressions are compiled as part of lowering to MIR, including checking for exhaustive patterns. Compilation follows a two-step process: first a match expression is compiled into a decision tree, then we lower the decision tree into MIR.


Ownership is enforced as follows: each expression stores its result in a virtual register, and MIR operates on these registers instead of variables. For each basic block we maintain a map of registers and their states, and blocks may inherit the state from parent blocks. When looking up the state of a register, we first check if the state is already known in the current block. If not, we look it up in its predecessors and store the resulting state in the current block. It's possible that two branches flow into a block, with each branch having a different state for a register. Consider this code for example:

let a = [10]

if true {
  let b = a
} else {


We can represent this code using the following graph:

| let a = [10] |
        v B            D
    +---------+     +-----+
    | if true | --> | bar | --+
    +---------+     +-----+   |       E
      |                       |   +-------+
      |                       +-> | after |
      v C                     |   +-------+
    +-----------+             |
    | let b = a | ----->------+

Now say we want to look up the state of the register storing a. If we follow the path A -> B -> C -> E, it's state is "moved" because of the expression let b = a. But if we follow the path A -> B -> D -> E the register is still available. The compiler handles this as follows: if we look up the state of a register and encounter multiple predecessor blocks, the state is a sum of the parent states. The state is calculated as follows:

  • If the state is "available" in all parents, the state is "available".
  • If the state is "moved" in all parents, the state is "moved".
  • If the state is "available" in one or more parents, but a different state in one or more other parents, the state is "maybe moved".

It's an error to use a register in state other than "available", and the compiler produces a corresponding error when encountering such code. These states are also used to determine how to drop values.

The compiler also has a fourth state: "partially moved". This is used when a moving method moves a field, in which case self is set to this state.

When traversing blocks to determine the register state, blocks already visited are skipped. If due to some sort of compiler bug we can't determine a state, the compiler terminates with a panic. This only happens in the event of a compiler bug.

When values are moved, the old virtual registers are marked as "moved". When exiting scopes, code is inserted to drop any registers not yet moved.

While the setup as described isn't that complex, implementing it correctly proved quite a challenge. For example, during the development we ran into bugs where explicit return expressions would drop registers not yet written to, because the compiler would allocate the register, process a sub expression containing a return, and not write a value until after processing that expression. To handle such cases, the compiler in various places allocates registers without storing them in the surrounding scope, then processes any sub expressions, and then adds the register to the surrounding scope.

Dropping values

When MIR is first generated, each drop results in a "drop" node. After performing the necessary analysis on MIR, these nodes are expanded into their final nodes. For example, a drop node acting on an owned value is transformed into three separate nodes:

  1. A check_refs node that ensures no references exist to the value to drop.
  2. A node that runs the dropper of the value.

The "dropper" is a generated method that recursively drops any fields, then runs the value's destructor (if any). Once the dropper finishes it releases the memory of the value.


MIR is also used to perform optimisations. For example, when sending a message to a process synchronously, the call is optimised into an instruction that doesn't allocate a future. MIR will also be used for future optimisations, such as inlining, devirtualisation, and dead code removal.

Code generation

Code generation is performed by lowering MIR into bytecode. This is done in a single pass over all the MIR. The result is a bytecode image for the VM to run.


As part of code generation, methods are stored in a class such that we can efficiently perform dynamic dispatch. Each class has a table for its method, with the size being a power of two. Each method is given a globally unique hash code based on its name. This means different methods with the same name share the same hash code, but this is OK because a type can only define a method with a given name once.

For each method we use the hash to calculate the ideal slot in the method table to place the method in. If a slot is already occupied we assign the method to the next available slot.

At runtime, dynamic dispatch is performed by taking the hash and hashing it into a method slot. Since method table sizes are a power of two, this is done using a bitwise AND instead of using the more expensive modulo operator. In other words, to get the slot we use the following:

slot = hash & (size - 1)

To handle conflicts we use linear probing, resulting in a reasonably efficient implementation of dynamic dispatch. For more information, refer to the article "Shenanigans With Hash Tables".