Logical Methods in Computer Science
Vol. 10(1:17)2014, pp. 1–52
www.lmcs-online.org

Submitted
Published

Aug. 27, 2012
Mar. 24, 2014

LINEAR USAGE OF STATE ∗
RASMUS EJLERS MØGELBERG a AND SAM STATON b
a

IT University of Copenhagen, Denmark
e-mail address: mogel@itu.dk

b

Radboud University Nijmegen, Netherlands
e-mail address: s.staton@cs.ru.nl
Abstract. We investigate the phenomenon that every monad is a linear state monad. We
do this by studying a fully-complete state-passing translation from an impure call-by-value
language to a new linear type theory: the enriched call-by-value calculus. The results
are not specific to store, but can be applied to any computational effect expressible using
algebraic operations, even to effects that are not usually thought of as stateful. There is a
bijective correspondence between generic effects in the source language and state access
operations in the enriched call-by-value calculus.
From the perspective of categorical models, the enriched call-by-value calculus suggests
a refinement of the traditional Kleisli models of effectful call-by-value languages. The new
models can be understood as enriched adjunctions.

1. Introduction
1.1. Informal motivation. The state-passing translation transforms a stateful program
into a pure function. As an illustration, consider the following ML program which uses a
single fixed memory cell l : int ref.
- fun f x = let val y = !l in l := x ; y end ;
val f = fn : int -> int
The state-passing translation transforms that program into the following pure function which
takes the state as an argument and returns the updated state as a result.
2012 ACM CCS: [Theory of computation]: Semantics and reasoning—Program semantics—
Denotational semantics / Categorical semantics; Logic—Linear logic / Type theory.
Key words and phrases: Linear type theory, monads, computational effects, categorical semantics, enriched
category theory, state passing translation.
∗
This article expands on a paper presented at the Fourth International Conference on Algebra and Coalgebra
in Computer Science (CALCO 2011).
a
Research supported by the Danish Agency for Science, Technology and Innovation.
b
Research supported by EPSRC Fellowship EP/E042414/1, ANR Projet CHOCO, the Isaac Newton Trust,
and ERC Projects ECSYM and QCLS.

l

LOGICAL METHODS
IN COMPUTER SCIENCE

c
DOI:10.2168/LMCS-10(1:17)2014

CC

R. E. Møgelberg and S. Staton
Creative Commons

2

R. E. MØGELBERG AND S. STATON

- fun f (x,s) = let val y = s val s’ = x in (y,s’) end ;
val f = fn : int * int -> int * int
The state-passing translation is straightforward if the program only uses a fixed, finite area
of memory of type S: an impure program of type A * B becomes a pure program of type
A × S → B × S.
To what extent does the state-passing translation apply to programs with other effects?
In this article we develop the idea that, from a semantic perspective, all effects can be
understood as state effects. Central to our treatment of state is the idea of linear usage: in
general, computations cannot copy the state and save it for later, nor can they discard the
state and insert a new one instead. In 1972, Strachey wrote [46]:
The state transformation produced by obeying a command is essentially
irreversible and it is, by the nature of the computers we use, impossible to
have more than one version of σ [the state] at any one time.
Historically, the importance of the linearity of state arose in the study of programs with
private store. In this setting, the naive state-passing translation does not preserve contextual
equivalence. For instance, the function ‘snapback’ takes a stateful computation f : A * B
and returns the computation that executes f but then snaps back to the original state:
def

snapback = λf : (A × S → B × S). λ(a, s) : A × S. hπ1 (f (a, s)), si : (A * B) → (A * B)
The snapback program does not arise as the state-passing translation of an impure program.
No impure program could tamper with the private store in the way that the snapback
program does. In other words, the state-passing translation is not fully complete. One
can use this fact to show that contextual equivalence is not preserved by the state-passing
translation. Sieber [44] insisted that every function be wrapped in snapback to obtain full
abstraction for his simple denotational model.
O’Hearn and Reynolds [32] resolved these difficulties with private store by moving to a
linear typing system. Linear usage of state can be expressed syntactically by considering a
stateful computation of type A * B as a linear map of type !A ⊗ S ( !B ⊗ S. The type of
states S must be used linearly, but the argument type A and the return type B can be used
arbitrarily. The snapback program violates these linear typing constraints.
This notation is reminiscent of Girard’s linear logic [13]. Our starting point is actually a
more refined calculus, the enriched effect calculus, which was developed by Egger, Møgelberg
and Simpson [7, 8, 9, 10] as a way of investigating linear usage of resources such as state.
All effects are state effects. In this paper we develop this linear typing discipline to show
that all effects can be understood as state effects and that there is always a fully-complete
linear-use state-passing translation. Our analysis applies even to effects that do not involve
the memory of a computer, like printing or coin-tossing. (For this reason we use the word
‘store’ to refer to memory related effects and ‘state’ for the general notion.) We now provide
two informal explanations of this general phenomenon.
A first informal explanation is that an inhabitant of the ‘state’ type S is an entire history
of the universe. The history of the universe certainly cannot be discarded or duplicated.
To be slightly more precise, if the effect in question is printing, then a ‘state’ is a string of
everything that has been printed so far.
A second informal explanation involves Jeffrey’s graphical notation [19]. Jeffrey noticed
that a naive graphical notation for composition of impure functions does not work, for it

LINEAR USAGE OF STATE

3

describes how functions depend on their arguments but it does not describe the order of
side effects:
l :=

x

'world'

print

'hello'

!l

print

To make the order of evaluation visible in the graphical notation, Jeffrey adds a special kind
of edge which must be treated linearly. This is what we call state.
l :=

x

print

'world'

!l

'hello'

print

Our contribution in this paper is foundational, but let us speculate briefly on possible applications. State plays an important role in many aspects of semantics, including operational
semantics and Hoare logic. The type !A ⊗ S can be thought of as a type of configurations
(program/state pairs) for a general abstract machine. This might pave the way for a general
framework for operational semantics, building on the ideas of Plotkin and Power [35].
We also note that variations on linear-use state-passing style are already used to
accommodate a broad class of effects within pure languages such as Clean [1] and Mercury [45].
1.2. The state-passing translation. The source language of our translation is an impure
functional language with product and function types:
σ, τ ::= 1 | σ × τ | σ * τ | . . .
We also consider sum types and unspecified base types. We adopt the call-by-value calling
convention because it is most natural one for effectful programs. For simplicity we choose
a fine-grain language in which the order of evaluation is totally explicit. The full syntax,
equational theory and framework for denotational semantics is in Section 3. Sum types are
treated in Section 6.
The types of the target language are essentially the minimal fragment of the enriched
effect calculus (EEC, [8, 7]) that is needed for the linear-use state-passing translation. To
enforce linear usage, EEC borrows some techniques and notation from linear logic (!, ⊗, and
() but in a careful way, enforced by the following grammar:
A, B ::= 1 | A × B | A ( B | . . .

value types

A, B ::= !A ⊗ B | S | . . .

computation types (underlined)

The full syntax, equational theory and framework for denotational semantics is in Section 2.
In Section 4 we provide a linear-use state passing translation from the source language
to the target language The translation on types takes a type τ of the source language to a
value type τ S of the target language, by induction on the structure of types:
def

1S = 1

def

(σ × τ )S = σ S × τ S

def

(σ * τ )S = !σ S ⊗ S ( !τ S ⊗ S

...

4

R. E. MØGELBERG AND S. STATON

Theorem 4.3 says that the linear-use state-passing translation is fully complete: it describes
a bijective correspondence between programs in the source language of type τ and programs
in the target language of type τ S . This means that contextual equivalence is preserved by
linear-use state-passing translation.
Enriched category theory. The constructions involved in the target language have an elegant
categorical formulation: they can be modelled in any enriched category with copowers. For
this reason we call the target language the enriched call-by-value calculus.
Recall that in categorical semantics a type A is interpreted as an object [[A]] of a category,
a context Γ is interpreted as an object [[Γ]] of a category, and a term-in-context Γ ` t : A is
interpreted as a morphism [[t]] : [[Γ]] → [[A]]. In the enriched call-by-value calculus, there are
two classes of type: computation types and value types. A categorical analysis must thus
involve two categories: a category V whose objects interpret value types, and a category C
whose objects interpret computation types. The structure of the types dictates the structure
that the categories V and C must have:
• For the product types (1, ×), V must have finite products in the categorical sense.
• For the tensor type (!A ⊗ B), V must act on C. This means that there is a functor
(−1 · −2 ) : V × C → C such that 1 · X ∼
= X and (A × B) · X ∼
= A · (B · X).
• For the linear function type (, the action (−1 · −2 ) must have a right adjoint in its first
argument.
The linear function type ( ensures that the space of morphisms X → Y forms an object
of V. So C is enriched in V, and the action structure provides copowers. There are
many examples of enriched models, including categories of algebras and Kleisli categories
(understood as ‘closed Freyd categories’ [27]). Full details are in Section 2.
In Section 5 we explain that the connection with Kleisli categories gives us a categorical
explanation of the linear-use state-passing translation. We use a semantic approach to
prove the full completeness of the translation. We show that the traditional models of
effectful call-by-value languages, using monads and Kleisli constructions, form a coreflective
subcategory of the enriched models, and that the state-passing translation is the unit of the
coreflection.
1.3. Relationship with monads. In an enriched situation, because S ( (−) is right
adjoint to !(−) ⊗ S, any computation type S induces a monad on value types:
S ( !(−) ⊗ S

(1.1)

We call this a linear-use state monad. In Section 3.2 we show that every monad arises as
a linear-use state monad. In brief, the argument is as follows. The right notion of monad
for programming languages is ‘strong monad T on a category V with finite products and
Kleisli exponentials’; for every such monad T , the Kleisli category C is V-enriched and has
copowers and is thus a model of the enriched call-by-value language. The object 1 in C
induces a linear-use state monad on V (1.1) which is isomorphic to the original monad T on
V.
For the simplest case, consider the following monad on the category of sets for storing a
single bit. If we let V = C = Set and let S = 2 (there are two states), then the linear-use
state monad is the usual state monad T X = 2 → (X × 2).

LINEAR USAGE OF STATE

5

In general, the linear-use state monad (1.1) arises from an adjunction that is parameterized in S. Atkey [3] has investigated monads that arise from parameterized adjunctions: in
fact Atkey’s parameterized monads are essentially the same as our enriched models. We
return to this in Section 9.
1.4. Algebraic theories and state access operations. In order to explain the general
connection between effects and state, we turn to the analysis of effects begun by Plotkin
and Power [38]. Categorically, a finitary monad on the category of sets is essentially the
same thing as an algebraic theory. Plotkin and Power proposed to use this connection to
investigate computational effects from the perspective of universal algebra.
The analysis of Plotkin and Power centres around the correspondence between algebraic
operations and generic effects. More precisely, for a monad T the following data are
equivalent:
(1) An algebraic operation: roughly, a functorial assignment of an n-ary function X n → X
to the carrier of each T -algebra T (X) → X;
(2) A generic effect: a function 1 → T (n).
For instance, consider the store monad for a single bit of memory (using the isomorphic
presentation T = ((−) × 2 × (−) × 2)). Each T -algebra ξ : T (X) → X supports an algebraic
operation ?X : X × X → X in a functorial way: let x ?X y = ξ(x, 0, y, 1). If we understand
elements of T -algebras as computations, then x ?X y is the computation that reads the bit
in memory and then branches to either x or y depending on what was read.
The corresponding generic effect deref : 1 → T (2) is given by deref() = (0, 0, 1, 1). It
can be thought of as the command that reads the bit and returns the contents of memory.
We have already explained (§1.3) that all monads can be understood as linear-use state
monads, T = (S ( (!(−) ⊗ S)). The data for algebraic operations and generic effects can
equivalently be given in terms of the following structure on the state object S:
(3) A state access operation: a morphism S → !n ⊗ S in the computation category C.
For instance, let V = C = Set and S = 2. This gives a state monad for a single bit of
memory, isomorphic to ((−) × 2 × (−) × 2). The state access operation corresponding to
reading the bit is simply the function read : 2 → 2 × 2 given by read(i) = (i, i), which reads
the store and returns the result along with the store.
In Section 8 we provide more examples and investigate a general framework for algebraic
presentations of theories of effects using algebraic operations, generic effects and state access
operations.
1.5. The enriched effect calculus. The work presented here grew out of work on the
enriched effect calculus by Egger, Møgelberg and Simpson (EEC, [7, 8, 9, 10]). The enriched
call-by-value calculus that we introduce in this paper is a fragment of EEC and our categorical
semantics is based on their work. Every model of EEC contains a monad, and one of the
observations of [9] (see also [7, Example 4.2]) was that this monad can always be represented
as a linear state monad. A special case of the embedding theorem [8, Theorem 4] shows
that given a strong monad T on a category V with finite products and Kleisli exponentials
we can embed V in a model of EEC preserving the monad and all other structure. This
gives the motto: every monad embeds in a linear state monad.
Since the enriched call-by-value calculus is a fragment of EEC (as opposed to full EEC),
we allow ourselves a broader class of models. In contrast to the earlier work on EEC,

6

R. E. MØGELBERG AND S. STATON

we do not include products among the computation types, since they are not needed in
the state-passing translation, and so in our models the category C does not need to have
products. This allows us to build models from Kleisli categories, which typically do not have
products, and this makes the relationship with monad models and closed Freyd categories
much more straightforward. In particular, in our setting every monad is a linear state
monad.
Nonetheless, in Section 10 we show that EEC is a conservative extension of the enriched
call-by-value calculus. This shows that there is a fully-complete linear-use state translation
into EEC. This result is further evidence that EEC is a promising calculus for reasoning
about linear usage of effects. The related papers [9, 10] show how the linear-use continuation
passing translation arises from a natural dual model construction on models of EEC, and
use this to prove a full completeness theorem similar to that proven here for the linear-use
state-passing translation. In fact, from the point of view of EEC the two translations
are surprisingly similar: the linear-use state-passing translation is essentially dual to the
linear-use continuation-passing translation. This observation goes back to the work on EEC
and indeed duality plays a key role in [9] (although the relationship with state wasn’t made
explicit there). We draw it out explicitly in Section 7.
Acknowledgements. We thank Alex Simpson for help and encouragement. Also thanks to
Lars Birkedal, Jeff Egger, Masahito Hasegawa, Shin-ya Katsumata and Paul Levy for helpful
discussions. Diagrams are typeset using the xymatrix package and Paul Taylor’s diagrams
package.
2. Enriched call-by-value: a calculus for enriched categories with copowers
The target language for the linear-use state translation is a new calculus called the enriched
call-by-value calculus (ECBV), that we now introduce. As we will explain, it is an internal
language for enriched categories with copowers.
The enriched call-by-value calculus is a fragment of the enriched effect calculus (EEC),
which was introduced by Egger et al. [8, 7] as a calculus for reasoning about linear usage
in computational effects. The types of ECBV can be understood as a fragment of linear
logic that is expressive enough to describe the linear state monad, S ( !(−) ⊗ S. We will
not dwell on the connection with linear logic here.
2.1. Type theory and equational theory. The enriched call-by-value calculus has two
collections of types: value types and computation types. We use α, β, . . . to range over a
set of value type constants, and α, β, . . . to range over a disjoint set of computation type
constants. We then use upright letters A, B, . . . to range over value types, and underlined
letters A, B, . . . to range over computation types, which are specified by the grammar below:
A, B ::= α | 1 | A × B | A ( B
A, B ::= α | !A ⊗ B .
Note that the construction !A ⊗ B is indivisible: the strings !α and α ⊗ β are not well-formed
types. The stratification of types means that one cannot chain function types: the string
α ( (β ( γ) is not well-formed.

LINEAR USAGE OF STATE

7

Readers familiar with Levy’s Call-by-Push-Value [26] or EEC [8] should note that there
are no type constructors F and U for shifting between value and computation types and
that computation types are not included in the value types. The only way to shift between
value types and computation types is by using tensor and function types. As we will see,
this is the essence of the state-passing translation.
The enriched call-by-value calculus has two basic typing judgements, written
Γ |− ` t: B

and

Γ |z : A ` t: B

(2.1)

In the both judgements, Γ is an assignment of value types to variables. In the first
judgement, B is a value type, and in the second judgement, both A and B need to be
computation types. The second judgement should be thought of as a judgement of linearity
in the variable z : A. These judgements are defined inductively by the typing rules in Figure 1,
which are a restriction of the rules of EEC [7] to this type structure. In the figure, ∆ is an
assignment of a computation type to a single variable, as in (2.1). The ideas behind the
term language for the function space A ( B and the tensor !A ⊗ B go back to the early work
on linear lambda calculus. In particular, the introduction rule for !A ⊗ B uses pairing, and
the elimination rule uses a pattern matching syntax.
2.2. Enriched call-by-value models. The categorical notion of model for ECBV involves
basic concepts from enriched category theory [21]. In summary, a model of the language
comprises two categories, V and C, interpreting the value and computation types respectively;
the function type A ( B provides an enrichment of C in V, and the universal property of
the tensor type !A ⊗ B is the copower, sometimes called tensor.
We now make this precise. Let us recall some rudiments. Following [18, 14], we
begin with actions of categories. Let V be a category with finite products (by which
we mean that it has a chosen terminal object and chosen binary products). Recall that
an action of V on a category C is a functor · : V × C → C together with two natural
isomorphisms, unit (1 · A) ∼
= A and associativity ((A × B) · C) ∼
= (A · (B · C)), that cohere
with the isomorphisms arising from the product structure of V in the following sense:
(A × 1) · D
∼
=

(1 × A) · D
∼
=



A · (1 · D)

%
/ A·D

∼
=



(A × (B × C)) · D

∼
=



1 · (A · D)
∼
=

((A × B) × C) · D
∼
=

∼
=

∼
=

/ A · ((B × C) · D)

∼
=

%
/ A·D

/ (A × B) · (C · D)


∼
=

∼
=

/ A · (B · (C · D))

(We underline objects of C to distinguish them from objects of V.)
An enrichment of a category C in V with copowers is determined by an action of V
on C such that each functor (− · B) : V → C has a right adjoint, C(B, −) : C → V. Then
A · B is called a copower, and C(B, C) is called enrichment. We write HomC (B, C) for the
usual hom-set of C to distinguish it from the enrichment.
Definition 2.1. An enriched call-by-value model (or simply enriched model ) is given by a
category V with finite products and a category C enriched in V with copowers.

8

R. E. MØGELBERG AND S. STATON

Types.
A, B ::= α | 1 | A × B | A ( B
A, B ::= α | !A ⊗ B .

Term formation.
Γ, x : A, Γ0 | − ` x : A
Γ |− ` t: A

Γ |z : A ` z : A

Γ |− ` u: B

Γ | − ` ht, ui : A × B
Γ |z : A ` t: B

Γ |− ` t: A

Γ | − ` t : A1 × A2
Γ | − ` πi (t) : Ai

Γ |− ` s: A ( B

Γ | − ` λz. t : A ( B
Γ |∆ ` u: B

Γ |− ` ?: 1

Γ |∆ ` t: A

Γ | ∆ ` s[t] : B
Γ | ∆ ` t : !A ⊗ B

Γ | ∆ ` !t ⊗ u : !A ⊗ B

Γ, x : A | z : B ` u : C

Γ | ∆ ` t to (!x ⊗ z). u : C

Equality. (We elide α-equivalence, reflexivity, symmetry, transitivity and congruence
laws.)
Γ |− ` t: 1

Γ | − ` t1 : A1

Γ |− ` t ≡ ?: 1

Γ | − ` t2 : A2

Γ | − ` πi (ht1 , t2 i) ≡ ti : Ai

Γ | − ` t : A1 × A2
Γ | − ` hπ1 (t), π2 (t)i ≡ t : A1 × A2
Γ |z : A ` t: B Γ |∆ ` u: A

Γ |− ` t: A ( B

Γ | ∆ ` (λz. t)[u] ≡ t[u/z] : B

Γ | − ` t ≡ λz. (t[z]) : A ( B

Γ | − ` t : A Γ | ∆ ` u : B Γ, x : A | z : B ` v : C
Γ | ∆ ` (!t ⊗ u) to (!x ⊗ z). v ≡ v[t/x, u/z] : C
Γ | ∆ ` t : !A ⊗ B Γ | y : !A ⊗ B ` u : C
Γ | ∆ ` t to (!x ⊗ z). u[!x ⊗ z/y] ≡ u[t/y] : C
Figure 1: The enriched call-by-value calculus

LINEAR USAGE OF STATE

9

In Section 2.3 we will illustrate the definition with some examples of enriched models.
First, let us clarify the semantics for ECBV in an enriched model. The interpretation is
similar to the semantics of EEC proposed by Egger et al. [8].
• A value type A is interpreted as an object [[A]] of V, and a computation type A is
interpreted as an object [[A]] of C, as follows. The interpretation is defined by induction
on the structure of types. First, for each value type constant α, an object [[α]] of V is
given, and for each computation type constant α an object [[α]] of C is given. The product
types are interpreted as products in V. The remaining type constructions are interpreted
def
def
using the enriched structure: we let [[!A ⊗ B]] = ([[A]] · [[B]]), and [[A ( B]] = C([[A]], [[B]]).
• A value context Γ = (x1 : A1 , . . . , xn : An ) is interpreted as a product [[A1 ]] × · · · × [[An ]].
in V. A computation context ∆ = (z : A) is interpreted as the object [[A]] in C.
• A judgement Γ | − ` t : A is interpreted as a morphism [[Γ]] → [[A]] in V, and a judgement
Γ | ∆ ` t : A is interpreted as a morphism [[Γ]] · [[∆]] → [[A]] in C. This definition is made
by induction on the structure of typing derivations, making use of the universal properties
of the interpretations of the types. For illustration, we consider the following two rules:
Γ |z : A ` t: B
Γ | ∆ ` t : !A ⊗ B Γ, x : A | z : B ` u : C
Γ | − ` λz. t : A ( B
Γ | ∆ ` t to (!x ⊗ z). u : C
In dealing with the linear lambda abstraction rule, the induction principle gives us an
interpretation [[t]] : [[Γ]] · [[A]] → [[B]] in C which we use to form [[λz. t]] : [[Γ]] → C([[A]], [[B]])
in V, using the natural bijection that defines the relationship between the copower and
the enrichment:
HomC (A · B, C) ∼
= HomV (A, C(B, C)).
For the pattern matching rule, we assume morphisms
[[t]] : [[Γ]] · [[∆]] → [[A]] · [[B]]

[[u]] : ([[Γ]] × [[A]]) · [[B]] → [[C]]

in C and use them to define [[t to (!x ⊗ z). u]] : [[Γ]] · [[∆]] → [[C]] as the following composite:
diag.

∼
=

[[Γ]]·[[t]]

[[Γ]] · [[∆]] −−−→ ([[Γ]] × [[Γ]]) · [[∆]] −
→ [[Γ]] · ([[Γ]] · [[∆]]) −−−−→ [[Γ]] · ([[A]] · [[B]])
∼
=

[[u]]

−
→ ([[Γ]] × [[A]]) · [[B]] −−→ [[C]]
Proposition 2.2. The interpretation of the enriched call-by-value calculus in an enriched
model (V, C) is sound:
(1) If Γ | − ` t ≡ u : A then [[t]] = [[u]] : [[Γ]] → [[A]] in V.
(2) If Γ | ∆ ` t ≡ u : A then [[t]] = [[u]] : [[Γ]] · [[∆]] → [[A]] in C.
Proof notes. This is proved by induction on the structure of the derivation of (≡). The
following substitution lemma is helpful:
If Γ | ∆ ` t : A and Γ | z : A ` u : B then u[t/z] is the following composite
morphism in C:
diag.

∼
=

[[Γ]]·[[t]]

[[u]]

[[Γ]] · [[∆]] −−−→ ([[Γ]] × [[Γ]]) · [[∆]] −
→ [[Γ]] · ([[Γ]] · [[∆]]) −−−−→ [[Γ]] · [[A]] −−→ [[B]]

10

R. E. MØGELBERG AND S. STATON

2.3. Examples of enriched models. We now list some examples of enriched models
(Definition 2.1).
(1) If V = Set then a V-enriched category is just a locally small category. The copower
A · B is the A-fold iterated coproduct of B, if it exists. The following three examples
are instances of this example.
(2) Let V = Set and let C be the category of monoids and homomorphisms. The copower
A · B, where A is a set and B is a monoid, can be described as a quotient of the free
monoid on the product of sets, (A × |B|)∗ /∼ . Here (A × |B|)∗ is the set of strings built
of pairs in (A × |B|), which is a monoid under concatenation with the empty string  as
unit. The equivalence relation (∼) is generated by (a, b).(a, b0 ) ∼ (a, b.b0 ) and  ∼ (a, ).
There is a similar description of the copower for any algebraic theory.
(3) We can modify Example (2) to make C the category of free monoids and monoid
homomorphisms. That is, the objects are monoids of the form B ∗ . In this situation, the
copower satisfies A · B ∗ = (A × B)∗ . In this example C is the Kleisli category for the
free monoid monad. We will revisit Kleisli categories in Section 3.2.
(4) Let V = C = Set, with C considered with the ordinary enrichment. The copower A · B
is the cartesian product of sets. This situation generalizes to the situation where V = C
is an arbitrary cartesian closed category.
(5) Let V be the category of ω-cpo’s and continuous functions, and let C be the category
of pointed ω-cpo’s and strict functions. The enrichment C(A, B) is the cpo of strict
functions under the pointwise order, and the copower A · B is the smash product A⊥ ⊗ B.
(6) In the next section we will investigate a model built from the syntax of the enriched
call-by-value calculus.
2.4. The syntactic enriched model. The types and terms of the enriched call-by-value
calculus form an enriched model which we call the syntactic enriched model.
Let V be the category whose objects are value types and where a morphism A → B
is a term in context x : A | − ` t : B modulo the equational theory (Figure 1) and modulo
renaming the free variable x. The identity morphism is the term x : A | − ` x : A, and
composition of morphisms
x : A|− ` t : B

y : B|− ` u : C

A −−−−−−−−→ B −−−−−−−−→ C
def

is given by substitution: u ◦ t = (x : A | − ` u[t/y ] : C). Since morphisms are actually
equivalence classes, the well-definedness of substitution depends on the following substitution
lemma
If x : A | − ` t ≡ t0 : B and y : B | − ` u ≡ u0 : C
0
then x : A | − ` u[t/y ] ≡ u0 [t /y ] : C
which is proved by induction on the derivation of u ≡ u0 .
The laws of associativity and identity for composition are immediate. For instance,
associativity amounts to the following syntactic identity:
If x : A | − ` t : B, y : B | − ` u : C and z : C | − ` v : D
t
then x : A | − ` (v[u/z ])[t/y ] ≡ v[u[ /y ]/z ] : D.
The category V has products, given by the product types. The equations at product
types are exactly what is needed to guarantee the universal properties.

LINEAR USAGE OF STATE

11

Let C be the category whose objects are computation types and where a morphism
A → B is a term in context − | z : A ` t : B modulo the equational theory and modulo
renaming the free variable z. Identities and composition are defined in a similar way to V.
The identity morphisms A → A are − | z : A ` z : A and composition is by substitution.
def
The action of V on C is given on objects by the tensor type: let A · B = !A ⊗ B. Given
morphisms
x : A|− ` t : A0

−|z : B ` u : B0

A −−−−−−−−→ A0 in V
and
B −−−−−−−−→ B0 in C
we define a morphism t · u : (A · B) → (A0 · B0 ) in C by

def
t · u = − | z 0 : !A ⊗ B ` z 0 to (!x ⊗ z). !t ⊗ u : !A0 ⊗ B0 .
Functoriality follows from the equational theory of ECBV. The unit and associativity
isomorphisms are straightforward to define. For example, associativity ((A × B) · C ∼
=
A · (B · C)) is given by exhibiting mutual inverses at all types:
− | z : !(A × B) ⊗ C ` z to (!x ⊗ z 0 ). !(π1 x) ⊗ (!(π2 x) ⊗ z 0 ) : !A ⊗ (!B ⊗ C)
− | z : !A ⊗ (!B ⊗ C) ` z to (!x ⊗ z 0 ). z 0 to (!y ⊗ z 00 ). !(x, y) ⊗ z 00 : !(A × B) ⊗ C
It follows from the equational theory of ECBV that these are isomorphisms and are natural
and coherent.
Finally, we discuss the enrichment of C in V. Given types A, B and C we describe a
natural bijection
HomC (A · B, C) ∼
= HomV (A, B ( C)
From left to right the bijection takes a computation term − | z : !A ⊗ B ` t : C to a value
term x : A | − ` λb. t[(!x⊗b)/z ] : B ( C. From right to left the bijection takes a value term
x : A | − ` u : B ( C to a computation term − | z : (!A ⊗ B) ` z to (!x ⊗ y). u[y] : C.
2.5. Universal property of the syntactic enriched model. The syntactic model described in Section 2.4 enjoys a universal property: it is an initial object in a category of
enriched models with structure preserving functors as morphisms. Given any other enriched
model, the unique morphism from the syntactic model is given by interpretation of syntax
in the model.
This semantic characterization of interpretation is standard in categorical semantics,
and it is useful for deriving syntactic results from semantics, as we shall see in Section 5.
However, we shall also see (Section 5.3) that we need to talk about morphisms of models
preserving structure only up to isomorphism, and the syntactic model is not initial with
respect to this collection of morphisms. Rather, interpretation defines a morphism which
is unique only up to unique isomorphism. In order to formulate this kind of initiality, we
need to be able to assess when two morphisms between enriched models are isomorphic.
Thus the enriched models form a groupoid-enriched category, i.e., a 2-category in which
each 2-cell is invertible. The idea of using groupoid-enriched categories of models has been
around for a long time (e.g. [6, §8]) and has been used in previous work on the enriched
effect calculus [8, 9].
A precise definition of the 2-category of enriched models Enr can be found in Appendix A.
Here we just state the initiality property of the syntactic model. First recall the following
definition from 2-category theory (see e.g. [22, §6]).

12

R. E. MØGELBERG AND S. STATON

Definition 2.3. Let K be a 2-category. An object 0 of K is bi-initial if for any object A
the hom-category K(0, A) is equivalent to the terminal category (i.e., the category with one
object and one morphism).
An equivalent way of stating bi-initiality is to say that for any other object A there
exists a 1-cell 0 → A which is unique up to unique 2-isomorphism.
The syntactic model of Section 2.4 is bi-initial in the category Enr, but in this paper
we are much more interested in a category of enriched models (V, C) with a specified state
object S in C (because of the relation to the state-passing translation), and so we formulate
bi-initiality with respect to a category Ecbv of these. Like all other structure, 1-cells of Ecbv
are only required to preserve the state objects up to isomorphism. (See Appendix A.2 for a
full definition). We write (Vecbv , Cecbv , S) for the enriched model obtained as in Section 2.4
from the syntax of the enriched call-by-value calculus extended with a special computation
type constant S.
Theorem 2.4. The model (Vecbv , Cecbv , S) is bi-initial in Ecbv. The unique morphisms
with domain (Vecbv , Cecbv , S) are given by interpretation of syntax into models.
3. Fine-grain call-by-value, a calculus for Kleisli models
The source language for our state-passing translation is a call-by-value language equipped
with an equational theory to be thought of as generated by some operational semantics, as
in [34]. We use a ‘fine grain’ call-by-value language, following Levy et al. [27, 26]. We use α
to range over type constants. The types are given by the grammar
σ, τ ::= α | 1 | σ × τ | σ * τ .
The function space * is a call-by-value one, which takes a value and produces a computation.
The fine-grain call-by-value calculus (FGCBV) has two typing judgements, one for values
and one for producers. These are written Γ `v V : σ and Γ `p M : σ. The latter should be
thought of as typing computations which produce values in the type judged but may also
perform side-effects along the way. In both judgements the variables of the contexts are to
be considered as placeholders for values. Typing rules along with equality rules are given in
Figure 2.
The call-by-value language is called ‘fine grain’ because the order of evaluation is explicit.
Notice that the string (f (x), g(y)) is not well-formed syntax: one must specify the order of
evaluation, for instance, like this:
f (x) to x0 . g(y) to y 0 . (x0 , y 0 ).
Translations from a ‘coarser grain’, more natural programming language are given by Levy
et al. ([27, §3], [26, §A.3.3]).
3.1. Interpretation in Kleisli models. The most natural way to interpret fine-grain
call-by-value is two have two categories V and C to interpret the judgements `v and `p ,
but to insist that the two categories have exactly the same objects, since in this language
there is only one class of types.

LINEAR USAGE OF STATE

13

Types.
σ, τ ::= α | 1 | σ × τ | σ * τ .

Term formation.
Γ `v V : σ1 × σ2
Γ, x : σ, Γ0 `v x : σ

Γ `v ? : 1

Γ `v V1 : σ1

Γ `v πi (V ) : σi

Γ `v V : σ

Γ `p M : σ

Γ `p return V : σ
Γ, x : σ `p N : τ

Γ `v V 2 : σ2

Γ `v hV1 , V2 i : σ1 × σ2
Γ, x : σ `p N : τ

Γ `p M to x. N : τ
Γ `v V : σ * τ

Γ `v λx. N : σ * τ

Γ `v W : σ

Γ `p V W : τ

Equality. (We elide α-equivalence, reflexivity, symmetry, transitivity and congruence
laws.)
Γ `v M : 1
Γ `v M ≡ ? : 1

Γ `v V1 : σ2

Γ `v V2 : σ2

Γ `v V : σ1 × σ2

Γ `v πi (hV1 , V2 i) ≡ Vi : σi

Γ, x : σ `p M : τ

Γ `v V : σ * τ

Γ `v V : σ

Γ `v (λx. M ) V ≡ M [V /x] : τ

Γ `v λx. (V x) ≡ V : σ * τ
Γ `v V : σ

Γ `p M : σ
Γ `p M to x. return x ≡ M : σ
Γ `p M : σ

Γ `v hπ1 (V ), π2 (V )i ≡ V : σ1 × σ2

Γ, x : σ `p N : τ

Γ `p return V to x. N ≡ N [V /x] : τ

Γ, x : σ `p N : τ

Γ, y : τ `p P : υ

Γ `p (M to x. N ) to y. P ≡ M to x. (N to y. P ) : υ
Figure 2: Fine-grain call-by-value.
Definition 3.1. An enriched Kleisli model is an enriched call-by-value model (V, C)
(Def. 2.1) together with an identity-on-objects functor J : V → C that strictly preserves
copowers, which means that J(A × B) = A · J(B) (naturally in A and B) and that the canonical isomorphisms induced by the product structure are the coherent unit and associativity
isomorphisms of the copowers:
1 · JA = J(1 × A) ∼
(A × B) · JC = J((A × B) × C) ∼
= JA
= J(A × (B × C)) = A · (B · JC).
We will sometimes say ‘Kleisli model’ for ‘enriched Kleisli model’. We use the name
‘Kleisli’ because this definition captures the situation where C is the Kleisli category for a
strong monad on V. The correspondence is explained in Section 3.2.

14

R. E. MØGELBERG AND S. STATON

Kleisli models have earlier been called ‘closed Freyd categories’ by Levy et al. [27].
Their original definition of closed Freyd category is based on premonoidal categories; the
relationship with actions of categories and Kleisli models is observed by Levy [26, B.10].
A semantics for FGCBV is given in a Kleisli model in a standard way.
• Each base type is given an interpretation [[α]] as an object of V. This interpretation is
extended to all types: [[1]] is the terminal object of V; [[σ × τ ]] is the product of [[σ]] and
def
[[τ ]]; and [[σ * τ ]] is defined using the enriched structure of C: [[σ * τ ]] = C([[σ]], [[τ ]]).
• A context Γ = (x1 : σ1 , . . . , xn : σn ) is interpreted in V as a product [[σ1 ]] × · · · × [[σn ]].
• A value type judgement Γ `v V : σ is interpreted as a morphism [[Γ]] → [[σ]] in V and a
producer type judgement Γ `p M : σ is interpreted as a morphism [[Γ]] → [[σ]] in C. This
is defined by induction on the structure of derivations, using the universal properties of
Kleisli models. For illustration we consider the following rule.
Γ `p M : σ Γ, x : σ `p N : τ
Γ `p M to x. N : τ
The induction hypothesis gives us two morphisms in C
[[M ]]

[[Γ]] −−→ [[σ]]

[[N ]]

[[Γ]] × [[σ]] −−→ [[τ ]]

and we use these to define a morphism in C that interprets (M to x. N ):
J(diag)

=

[[Γ]]·[[M ]]

=

[[N ]]

[[Γ]] −−−−→ [[Γ]] × [[Γ]] −
→ [[Γ]] · [[Γ]] −−−−−→ [[Γ]] · [[σ]] −
→ [[Γ]] × [[σ]] −−→ [[τ ]] .
As another example [[return V ]] = J([[V ]]).
This defines a sound and complete notion of model for FGCBV.
Proposition 3.2 ([27], Prop. 5.1). The interpretation of the fine-grain call-by-value calculus
in a Kleisli model is sound:
(1) If Γ `v V ≡ W : σ then [[V ]] = [[W ]] : [[Γ]] → [[σ]] in V.
(2) If Γ `p M ≡ N : σ then [[M ]] = [[N ]] : [[Γ]] → [[σ]] in C.
3.2. Relationship with monads. We now explain the connection between enriched Kleisli
models and Kleisli categories for a monad. For more detail, see the paper by Levy et al. [27].
From the syntactic point of view, the fine-grain call-by-value language can be thought
of as a variant of Moggi’s λC [31]: the type construction (1 * (−)) is a monad.
From the semantic point of view, recall that a λC -model [31] is a category with finite
products and a strong monad with Kleisli exponentials. We now explain these conditions.
Let V be a category with finite products, and let (T, η, µ) be a monad on V. Let C be
the Kleisli category for T : the objects of C are the objects of V and a morphism A → B in
C is a morphism A → T (B) in V. There is an identity-on-objects functor J : V → C which
takes a morphism f : A → B in V to the morphism (ηB · f ) : A → B in C.
A strength for a monad T is usually expressed as a family of morphisms A × T (B) →
T (A × B) that respect the structure of the monad. In fact, a monad is strong if and only if
there is an action of V on C and the identity-on-objects functor J : V → C preserves it. The
strength is needed to take a morphism f : B → B 0 in C to a morphism (A · f ) : A · B → A · B 0
in C.
The requirement of Kleisli exponentials is normally expressed as the requirement that
for all A and B, the hom-functor HomV ((−) × A, T B) : Vop → Set is representable. To

LINEAR USAGE OF STATE

15

endow V with Kleisli exponentials is to give a right adjoint for the action, i.e. an enrichment
of C in V.
Conversely, every enriched Kleisli model (V, C, J) induces a strong monad on V with
def
Kleisli exponentials. The monad is defined using the closed structure: T (A) = C(1, A).
The Kleisli category for this monad is isomorphic to C. On the other hand, if we begin with
a monad, build the Kleisli category and then take the monad C(1, A), we recover a monad
that is isomorphic to the one that we started with. In this sense, enriched Kleisli models and
λC -models are equivalent. Note that they are not exactly the same, for although the hom-set
HomV (A, C(1, B)) is in bijection with HomC (A, B), the sets are typically not identical.
3.3. The syntactic Kleisli model. The types and terms of the fine-grain call-by-value
calculus form a syntactic model. We employ the same kinds of technique as for enriched
call-by-value in Section 2.4.
• The objects of both V and C are the types of FGCBV.
• A morphism σ → τ in V is a value judgement x : σ `v V : τ modulo the equational
theory ≡ (Figure 2) and modulo renaming the free variable x. Identities are variables and
composition is by substitution.
• A morphism σ → τ in C is a computation judgement x : σ `p M : τ modulo the equational
theory ≡ and renaming the free variable x. The identity σ → σ is x : σ `p return x : σ.
Composition is not by substitution, since one cannot substitute a producer term for a
variable. Rather, the composite of
x : σ`p M : τ

y : τ `p N : υ

σ −−−−−−−→ τ −−−−−−−→ υ
in C is x : σ `p M to y. N : υ.
• The product structure in V is given by the product types, projections and pairing.
def
• The action of V on C is given on objects by the binary product types: let σ · τ = σ × τ .
y : τ `p M : τ 0

x : σ`v V : σ 0

On morphisms, given A −−−−−−−→ σ 0 in V and τ −−−−−−−→ τ 0 in C, we define
V ·M

def

(σ · τ −−−→ σ 0 · τ 0 ) = z : σ × τ `p M [π2 (z)/y ] to y 0 . return hV [π1 (z)/x ], y 0 i : σ 0 × τ 0
def

• The enrichment is given by C(σ, τ ) = (σ * τ ).
x : σ`v V : τ

• The identity-on-objects functor J : V → C takes a morphism σ −−−−−−−→ τ in V to the
x : σ`p return V : τ
morphism σ −−−−−−−−−−−→ τ in C.
We write (Vfgcbv , Cfgcbv , J) for the syntactic Kleisli model.
3.4. Universal property of the syntactic Kleisli model. Appendix A.3 defines the
2-category Kleisli of Kleisli models. As was the case for Ecbv, 1-cells are only required to
preserve structure up to isomorphism.
Theorem 3.3. The syntactic Kleisli model (Vfgcbv , Cfgcbv , J) is bi-initial in Kleisli. The
unique morphisms with domain (Vfgcbv , Cfgcbv , J) are given by interpretation of syntax into
models.

16

R. E. MØGELBERG AND S. STATON

4. The linear-use state-passing translation
This section defines the linear-use state-passing translation from the fine-grain call-by-value
calculus to the enriched call-by-value calculus, and states the main syntactic results of this
paper: fullness on types and full completeness. Together these assert that the linear-use
state-passing translation is an equivalence of languages.
We now fix a computation type S of ECBV. For now, it can be an arbitrary computation
type; later we will make it a distinguished basic type to achieve a full completeness result.
We will describe a translation from FGCBV to ECBV. When S is thought of as a type of
states, then this translation reads as a state-passing translation.
We translate FGCBV types σ to ECBV value types σ S :
def

αS = α

def

(σ × τ )S = σ S × τ S

def

1S = 1

def

(σ * τ )S = !(σ S ) ⊗ S ( !(τ S ) ⊗ S

We extend this translation to type contexts, taking an FGCBV type context Γ to an ECBV
type context ΓS .
The translation on terms is syntax-directed. We pick a variable s, completely fresh.
The translation takes an FGCBV value type judgement Γ `v V : σ to an ECBV judgement
ΓS | − ` V S : σ S , and an FGCBV producer judgement Γ `p M : σ to an ECBV judgement
S
ΓS | s : S ` Ms : !(σ S ) ⊗ S. The translation is defined as follows.
def

xS = x

def

?S = ?

def

hV, W iS = hV S, W S i
def

(return V )Ss = !(V S ) ⊗ s

def

(π1 (V ))S = π1 (V S )

def

(π2 (V ))S = π2 (V S )

def

(M to x. N )sS = MsS to (!x ⊗ s). NsS
def

def

(λx. N )S = λz. z to (!x ⊗ s). NsS
(V W )sS = V S [!(W S ) ⊗ s]
In the case for λ-abstraction, the z is chosen to be completely fresh.
The translation respects types. For instance
Γ, x : σ `p N : τ
Γ `v λx. N : σ * τ

ΓS | z : !σ S ⊗ S ` z : !σ S ⊗ S
becomes

ΓS , x : σ S | s : S ` NsS : !τ S ⊗ S

ΓS | z : !σ S ⊗ S ` z to (!x ⊗ s). NsS : !τ S ⊗ S
ΓS | − ` λz. z to (!x ⊗ s). NsS : !σ S ⊗ S ( !τ S ⊗ S

Theorem 4.1. The linear-use state-passing translation is sound:
(1) If Γ `v V ≡ W : σ then ΓS | − ` V S ≡ W S : σ S .
S
S
(2) If Γ `p M ≡ N : σ then ΓS | s : S ` Ms ≡ Ns : σ S .
This result can be proved by induction on the structure of equality (≡) derivations, but
it can also be derived semantically as we shall see in Section 5.4.
4.1. Full completeness. We now state our main theorems: fullness on types and full
completeness on terms. To state fullness on types we need to talk about isomorphism of
types in ECBV. This can be defined in the usual way: for value types, an isomorphism A ∼
=B
is given by two judgements, x : A | − ` t : B and y : B | − ` u : A, such that u[t/y] ≡ x and
t[u/x] ≡ y. For computation types, A ∼
= B is witnessed by closed terms of type A ( B,
B ( A composing in both directions to identities. We note the following type isomorphisms,
inherited from the enriched effect calculus [8, §3]:
A ∼
!A ⊗ (!B ⊗ C) ∼
(4.1)
= !1 ⊗ A
= !(A × B) ⊗ C

LINEAR USAGE OF STATE

17

Theorem 4.2 (Fullness on types). Let A be a value type of ECBV formed using no other
computation type constants than S. Then there exists an FGCBV type σ such that σ S ∼
= A.
Proof. By induction on the structure of types. The interesting case A ( B uses the fact
that any computation type not using any α other than S is isomorphic to one of the form
!C ⊗ S, which follows from the isomorphisms (4.1).
We now state our main syntactic result.
Theorem 4.3 (Full completeness).
(1) If Γ `v V, W : σ and ΓS | − ` V S ≡ W S : σ S then Γ `v V ≡ W : σ.
S
S
(2) If Γ `p M, N : σ and ΓS | s : S ` Ms ≡ Ns : !σ S ⊗ S then Γ `p M ≡ N : σ.
(3) For any ΓS | − ` t : σ S there is a term Γ `v V : σ such that ΓS | − ` t ≡ V S : σ S .
(4) For any ΓS | s : S ` t : !(σ S ) ⊗ S there is a producer term Γ `p M : σ such that ΓS | s : S `
t ≡ M S : !(σ S ) ⊗ S.
In Section 5.4 we sketch a semantic proof of Theorems 4.2 and 4.3.
5. A semantic proof of full completeness
In this section we present two constructions on models. The first (§5.1) constructs a Kleisli
model (Def. 3.1) from an enriched model (Def. 2.1) with a specified state object. The second
(§5.2) constructs an enriched model with a state object from a given Kleisli model. The
state-passing translation arises from the first of these constructions. These two constructions
form a bi-adjunction exhibiting the category of Kleisli models as a coreflective subcategory of
the category of enriched models with chosen state objects (§5.3). In Section 5.4 we shall see
how to use these facts to explain full completeness of the linear-use state-passing translation
(Theorem 4.3).
5.1. From enriched models with state to Kleisli models. Given an enriched call-byvalue model (V, C) with a state object S in C, we can form an enriched Kleisli model
def
Kl(V, C, S) = (V, KlS , JS ), where the category KlS has the same objects as V and
hom-sets
def
HomKlS (A, B) = HomC (A · S, B · S)
Composition in KlS is just composition as in C. (This is an isomorphic presentation of the
Kleisli category for the monad C(S, − · S) on V.) The functor JS is the identity on objects
and maps f : A → B to f · S.
Lemma 5.1. For any enriched call-by-value model with state (V, C, S) the data (V, KlS , JS )
defines an enriched Kleisli model.
Proof. The action (−1 ) ·Kl (−2 ) : V × KlS → KlS is defined on objects as A ·Kl B = A × B.
On morphisms it maps f : A → A0 , g : B · S → B 0 · S to the composite
∼
=

f ·g

∼
=

→ A · (B · S) −−→ A0 · (B 0 · S) −
(A × B) · S −
→ (A0 × B 0 ) · S
which is easily seen to be functorial.
def
The right adjoint to (−) ·Kl A is KlS (A, −) = C(A · S, (−) · S).

18

R. E. MØGELBERG AND S. STATON

The construction Kl described above extends to a 2-functor Kl : Ecbv → Kleisli from
the 2-category of enriched models to the 2-category of Kleisli models. See Appendix C.2 for
details.
5.2. From Kleisli models to enriched models with state. Any Kleisli model is trivially
an enriched model, so for the opposite construction we just need to pick a state object in
def
a Kleisli model. We define St(V, C, J) = (V, C, 1), where 1 is the terminal object of V
considered as an object of C. This definition extends to a 2-functor St : Kleisli → Ecbv, as
shown in Appendix C.
The motivation for this definition is that, as we now show, the 2-category Ecbv can
be seen as a 2-category of enriched adjunctions, and the 2-functor St can be seen as an
inclusion of Kleisli adjunctions into Ecbv.
Let (V, C) be an enriched model. By an enriched adjunction we mean an adjunction
F a U : C → V equipped with a natural coherent isomorphism F (A × B) ∼
= A · F (B). When
V is cartesian closed, this is equivalent to the usual definition, i.e. a natural isomorphism
C(F (−1 ), −2 ) ∼
= V(−1 , U (−2 )) in V (see e.g. [20]).
Any choice of state object gives an enriched adjunction, since (− · S) is left adjoint to
C(S, −) : C → V. The following proposition (first noted for EEC, [8, Proof of Thm. 4], [7])
shows that every enriched adjunction arises in this way:
Proposition 5.2 ([7]). Let (V, C) be an enriched model. If F a U : C → V is an enriched
adjunction then it is naturally isomorphic to the enriched adjunction induced by F (1).
So enriched adjunctions correspond essentially bijectively to state objects. In particular
the state object corresponding to a Kleisli model is 1. Monads induced by state objects can
be described in ECBV as S ( !(−) ⊗ S. By the correspondence between Kleisli models and
strong monads we arrive at the slogan: Every strong monad is a linear-use state monad.
More directly, the slogan can be derived from the isomorphism KlT (1, A × 1) ∼
= T (A), which
holds for the Kleisli category KlT of any strong monad T .
(Remark: if T is a strong monad on V then, for any object S of KlT , the linear-use
state monad KlT (S, (−) · S) on V is also known as the application of the state monad
transformer to T , as in [28].)
5.3. A coreflection. The constructions Kl and St form a bi-adjunction between the 2categories of Kleisli models and of enriched models with state. One intuition for this is that
it expresses the minimality property of Kleisli resolutions of monads.
Theorem 5.3. The 2-functor St : Kleisli → Ecbv is left biadjoint to Kl, i.e., for any pair
of objects (V, C, J) and (V0 , C0 , S) of Kleisli and Ecbv respectively, there is an equivalence
of categories
Ecbv(St(V, C, J), (V0 , C0 , S)) ' Kleisli((V, C, J), Kl(V0 , C0 , S))
natural in (V, C, J) and (V0 , C0 , S). Moreover, the unit of the adjunction η : id Ecbv → Kl◦St
is an isomorphism.
See Appendix C for a proof of Theorem 5.3. Since left bi-adjoints preserve bi-initial
objects we get the following connection between the syntactic enriched model (Vecbv , Cecbv , S)
and the syntactic Kleisli model (Vfgcbv , Cfgcbv , J).

LINEAR USAGE OF STATE

19

Corollary 5.4. (Vecbv , Cecbv , S) and St(Vfgcbv , Cfgcbv , J) are equivalent as objects of Ecbv.
Since Kl(St(V, C, J)) = (V, Kl1 , J1 ) the unit of the adjunction can be described as the
pair (id V , G : C → Kl1 ) where G(f : A → B) is the composite
∼
=

f

∼
=

A·1−
→A−
→B−
→B·1
using the isomorphism J(π) : A · 1 = A × 1 → A. The pair (id V , G) preserves the enrichment
only up to isomorphism, and this is our main motivation for using 2-categories of models
(see also the discussion in Section 2.5).
5.4. A semantic explanation of the state-passing translation. The linear-use statepassing translation is essentially the interpretation of fine-grain call-by-value into the model
obtained by applying the construction Kl of Section 5 to the syntactic enriched call-by-value
model of Section 2.4. In this model judgements Γ `v V : σ and Γ `p M : σ are interpreted as
judgements of the form
Q S
Q
x:
Γ | − ` [[V ]] : σ S
− | x : !( ΓS ) ⊗ S ` [[M ]] : σ S
Q
respectively, where ΓS is the product of all the types appearing in ΓS .
Lemma 5.5. Let (Vecbv , Cecbv , S) be the syntactic enriched Kleisli model of Section 2.4.
The interpretation of FGCBV into Kl(Vecbv , Cecbv , S) models a type σ as σ S . Let Γ =
x1 : σ1 . . . xn : σn be a context of FGCBV and let Γ `v V : τ and Γ `p M : τ be judgements of
FGCBV. Then V and M are interpreted as the equivalence classes of the terms
Q
x : ( ΓS ) | − ` V S [π1 x...πn x/x1 ...xn ] : τ S
Q
− | x : !( ΓS ) ⊗ S ` x to (!z ⊗ s). (M S [π1 z...πn z/x1 ...xn ]) : τ S
Soundness of the state-passing translation (Theorem 4.1) follows immediately from
Lemma 5.5. Fullness on types and full completeness (Theorems 4.2 and 4.3) are also
consequences.
Proof of Theorems 4.2 and 4.3. By Theorem 3.3 and Lemma 5.5 the state-passing translation is the unique (up to isomorphism) 1-cell
(F, G) : (Vfgcbv , Cfgcbv , J) → Kl(Vecbv , Cecbv , S)
in Kleisli. It suffices to show that this is an equivalence in Kleisli, because this implies that
F and G are both equivalences of categories, in particular they are essentially full on objects
(proving Theorem 4.2) and full and faithful (proving Theorem 4.3).
By initiality (F, G) must be isomorphic to the composite
η

'

(Vfgcbv , Cfgcbv , J) −
→ Kl(St(Vfgcbv , Cfgcbv , J)) −
→ Kl(Vecbv , Cecbv , S)
of the unit of the adjunction (which is an isomorphism by Theorem 5.3) and Kl applied to
the equivalence of Corollary 5.4. Since this composite is an equivalence, so is (F, G).

20

R. E. MØGELBERG AND S. STATON

6. Sums
It is routine to add sum types to the languages considered in Section 2 and 3, and the
state-passing translation extends straightforwardly. We now summarize the details.
6.1. Sums in the enriched call-by-value calculus. We add sums to the enriched callby-value calculus, for both value and computation types. The required modifications to the
calculus are given in Figure 3. The resulting calculus is still a fragment of the enriched effect
calculus [7]. We now extend the notion of model to accommodate the sums.
Terminology. Recall that a distributive category is a category with finite products and
coproducts such that the canonical morphisms 0 → A × 0 and ((A × B) + (A × C)) →
(A × (B + C)) are isomorphisms.
If a distributive category V has an action (·) on a category C with finite coproducts
(0, ⊕), then we say that the situation is distributive if the four canonical morphisms 0 → A · 0,
(A · B) ⊕ (A · C) → A · (B ⊕ C), 0 → 0·A and (A·C)⊕(B·C) → (A+B)·C are isomorphisms.
If the action is enriched, i.e. each functor (− · A) : V → C has a right adjoint, then this
definition of distributive coproducts amounts to the usual notion of enriched coproducts.
(Note that when the action is enriched then the third and fourth canonical morphisms are
necessarily isomorphisms since left adjoints preserve colimits.)
Definition 6.1. A distributive enriched model is given by a distributive category V and a
category C enriched in V with copowers and enriched coproducts.
It is straightforward to extend the sound interpretation of the enriched call-by-value
calculus in enriched models (§2.2) to a sound interpretation of enriched call-by-value calculus
with sums in distributive enriched models.
Of the examples in Section 2.3, (1)–(5) are distributive. The syntactic model of the
version of the calculus with sums is a distributive enriched model, and it is bi-initial for the
suitable generalization of morphism.
6.2. Sums in the fine-grain call-by-value calculus. It is equally straightforward to add
sums to the fine-grain call-by-value language. This is summarized in Figure 4.
We only include constructors/destructors as value terms, but from these we can derive
constructors/destructors for producer terms, as follows.
def

imagep (M ) = M to x. return ?(x)

def

inpi (M ) = M to x. return ini (x)

casep M of (in1 (x1 ).N1 |in2 (x2 ).N2 )
def

= M to z. (case z of (in1 (x1 ).λw. N1 |in2 (x2 ).λw. N2 ))(?)

where w, z are fresh.
These constructions have derived typing rules
Γ ` p Mi : σ i
Γ `p M : 0
Γ `p imagep (M ) : σ

Γ `p inpi (M ) : σ1 + σ2

(i = 1, 2)
(6.1)

p

Γ ` M : σ1 + σ2
p

p

p

Γ, xi : σi ` Ni : τ (i = 1, 2)

Γ ` case M of (in1 (x1 ).N1 |in2 (x2 ).N2 ) : τ

LINEAR USAGE OF STATE

21

Types.
A, B ::= α | 1 | A × B | A ( B | 0 | A + B
A, B ::= α | !A ⊗ B | 0 | A ⊕ B .

Term formation. The following rules are in addition to the rules in Figure 1.
Γ |− ` t: 0

Γ |∆ ` t: 0

Γ | − ` ?(t) : A

Γ | ∆ ` ?(t) : A

Γ | − ` t : Ai
Γ | − ` ini (t) : A1 + A2
Γ | − ` t : A1 + A2

(i = 1, 2)

Γ | ∆ ` t : Ai
Γ | ∆ ` ini (t) : A1 ⊕ A2

Γ, x1 : A1 | − ` u1 : C

(i = 1, 2)

Γ, x2 : A2 | − ` u2 : C

Γ | − ` case t of (in1 (x1 ). u1 |in2 (x2 ). u2 ) : C
Γ | ∆ ` s : A1 ⊕ A2

Γ | x1 : A1 ` t1 : C

Γ | x2 : A2 ` t2 : C

Γ | ∆ ` case s of (in1 (x1 ). t1 |in2 (x2 ). t2 ) : C

Equality. The following rules are in addition to the rules in Figure 1.
Γ |− ` t: 0

Γ, x : 0 | − ` u : A

Γ |∆ ` t: 0

t

Γ | ∆ ` ?(t) ≡ u[t/x ] : A

Γ | − ` ?(t) ≡ u[ /x ] : A
Γ | − ` t : Ai

Γ |x: 0 ` u: A

Γ, x1 : A1 | − ` u1 : B

Γ, x2 : A2 | − ` u2 : B

Γ | − ` case (ini (t)) of (in1 (x1 ). u1 |in2 (x2 ). u2 ) ≡ ui [t/xi ] : B
Γ | ∆ ` t : Ai

Γ | x1 : A1 ` u1 : B

Γ | x2 : A2 ` u2 : B

Γ | ∆ ` case (ini (t)) of (in1 (x1 ). u1 |in2 (x2 ). u2 ) ≡ ui [t/xi ] : B
Γ | − ` t : A1 + A2

(i = 1, 2)

(i = 1, 2)

Γ, z : A1 + A2 | − ` u : B

Γ | − ` u[t/z ] ≡ case t of (in1 (x1 ). u[in1 (x1 )/z ]|in2 (x2 ). u[in2 (x2 )/z ]) : B
Γ | ∆ ` t : A1 ⊕ A2

Γ | z : A1 ⊕ A2 ` u : B

Γ | ∆ ` u[t/z ] ≡ case t of (in1 (x1 ). u[in1 (x1 )/z ]|in2 (x2 ). u[in2 (x2 )/z ]) : B
Figure 3: Additional rules for sum types in Enriched Call-by-Value

22

R. E. MØGELBERG AND S. STATON

Types.
σ ::= α | 1 | σ × σ | σ * σ | 0 | σ + σ

Term formation. The following rules are in addition to the rules in Figure 2.
Γ `v V : 0

Γ `v V : σ1

Γ `v V : σ2

Γ `v ?(V ) : σ

Γ `v in1 (V ) : σ1 + σ2

Γ `v in2 (V ) : σ1 + σ2

Γ `v V : σ1 + σ2

Γ, x1 : σ1 `v W1 : τ

Γ, x1 : σ2 `v W2 : τ

Γ `v case V of (in1 (x1 ).W1 |in2 (x2 ).W2 ) : τ

Equality. The following rules are in addition to the rules in Figure 2.
Γ `v V : 0

Γ, x : 0 `v W : σ

Γ `v ?(V ) ≡ W [V /x ] : σ
Γ `v V : σi

Γ, x1 : σ1 `v W1 : τ

Γ, x1 : σ2 `v W2 : τ

Γ `v case ini (V ) of (in1 (x1 ).W1 |in2 (x2 ).W2 ) ≡ Wi [V /xi ] : τ
Γ `v V : σ1 + σ2

(i = 1, 2)

Γ, z : σ1 + σ2 `v W : τ

Γ `v W [V /z ] ≡ case V of (in1 (x1 ).W [in1 (x1 )/z ]|in2 (x2 ).W [in2 (x2 )/z ]) : τ
Figure 4: Additional rules for sum types in Fine-Grain Call-by-Value.
For example:
Γ, x1 : σ1 `p N1 : τ

Γ, x2 : σ2 `p N2 : τ

Γ, x1 : σ1 `v λw. N1 : 1 * τ

Γ, x2 : σ2 `v λw. N2 : 1 * τ

Γ, z `v case z of (in1 (x1 ).λw. N1 |in2 (x2 ).λw. N2 ) : 1 * τ
Γ `p M : σ1 + σ2

Γ, z `v ? : 1

Γ, z : σ1 + σ2 `p (case z of (in1 (x1 ).λw. N1 |in2 (x2 ).λw. N2 )) (?) : τ
Γ `p casep M of (in1 (x1 ).N1 |in2 (x2 ).N2 ) : τ

We now refine the notion of Kleisli model (Def. 3.1) to accommodate the sums.
Definition 6.2. A distributive enriched Kleisli model (distributive Kleisli model for short)
is a distributive enriched model (Def 6.1) together with an identity-on-objects functor
J : V → C that strictly preserves copowers and coproducts.
Note that in any Kleisli model, J will preserve coproducts because it has a right adjoint,
C(1, −). We insist moreover that it preserves coproducts strictly.
Note that a distributive Kleisli model is what Power [41, Def. 36] calls a distributive
closed Freyd category.

LINEAR USAGE OF STATE

23

It is straightforward to extend our interpretation of the fine-grain call-by-value calculus
in Kleisli models (§3.1) to an interpretation of the calculus with sums in distributive Kleisli
models. The interpretation is sound and there is a syntactic model.
In Section 3.2 we discussed the equivalence between enriched Kleisli models and strong
monads with Kleisli exponentials. This equivalence extends to an equivalence between
distributive enriched Kleisli models, and strong monads on distributive categories with Kleisli
exponentials. Likewise, the constructions St and Kl of Section 5 extend to constructions
deriving distributive enriched models from distributive Kleisli models and vice versa, and an
extension of Theorem 5.3 states that St and Kl exhibit a 2-category of distributive Kleisli
models as a coreflective subcategory of a 2-category of distributive enriched models.
6.3. Sums and the state-passing translation. It is straightforward to adapt the statepassing translation to accommodate sums. Recall that each type σ of FGCBV is translated
to a value type σ S of ECBV. We set
def

0S = 0

def

(σ + τ )S = σ S + τ S .

We extend this translation to type contexts, taking an FGCBV type context Γ to an ECBV
type context ΓS .
Recall that the translation on terms takes an FGCBV value type judgement Γ `v V : σ to
an ECBV judgement ΓS | − ` V S : σ S , and takes an FGCBV producer judgement Γ `p M : σ
S
to an ECBV judgement ΓS | s : S ` Ms : !(σ S ) ⊗ S. We extend the translation in Section 4
straightforwardly, as follows:
def

?(V )S = ?(V S )

def

in1 (V )S = in1 (V S )
def

def

in2 (V )S = in2 (V S )
S

S

case V of (in1 (x1 ).W1 |in2 (x2 ).W2 )S = case V S of (in1 (x1 ). W1 |in2 (x2 ). W2 )
The translation remains sound. The full definability results (Theorems 4.2 and 4.3) continue
to hold in the presence of sums.
7. Remarks on the linear-use continuation-passing translation
We now briefly emphasise that the linear-use continuation-passing translation arises as
a formal dual of the linear-use state-passing translation. This is not a new observation:
Hasegawa noticed it in the context of classical linear logic ([15, §8],[33]) and indeed it
informed the earlier work on the enriched effect calculus.
The linear-use continuation-passing translation was first elaborated by Berdine, O’Hearn,
Reddy and Thielecke [4], but our main reference is the more recent work by Egger, Møgelberg
and Simpson [9, 10]. They showed that the linear-use continuation-passing translation can be
extended to an involutive translation of the enriched effect calculus to itself, and derived a full
completeness result from this. That work, in turn, stems from Hasegawa’s full completeness
result [15] for a linear-use continuation-passing translation into dual intuitionistic / linear
logic.
Following [9], our development is fuelled by the following categorical observation. If
a category C is enriched in V with copowers, then we can form the dual category Cop
which is also enriched in V, but now with powers instead of copowers. (Recall that an
enriched category C has powers Y X if the functor HomV (X, C(−, Y )) : Cop → Set is

24

R. E. MØGELBERG AND S. STATON

Types.
A, B ::= α | 1 | A × B | A ( B
A, B ::= α | A → B .

Term formation.
Γ, x : A, Γ0 | − ` x : A
Γ |− ` t: A

Γ |z : A ` z : A

Γ |− ` u: B

Γ | − ` ht, ui : A × B
Γ |z : A ` t: B

Γ | − ` t : A1 × A2
Γ | − ` πi (t) : Ai

Γ |− ` s: A ( B

Γ | − ` λz. t : A ( B
Γ, x : A | ∆ ` t : B

Γ |− ` ?: 1

Γ |∆ ` t: A

Γ | ∆ ` s[t] : B
Γ |∆ ` t: A → B

Γ | ∆ ` λx. t : A → B

Γ, x : A | − ` u : A

Γ |∆ ` t u: B

Equations: Equations of Figure 1 but with equations for tensor types replaced by:
Γ, x : A | ∆ ` t : B Γ | − ` u : A

Γ |∆ ` t: A → B

Γ | ∆ ` (λx. t) u ≡ t[u/x] : B

Γ, x : A | ∆ ` t ≡ λx. (t x) : A → B

Figure 5: A CPS variant of the enriched call-by-value calculus.
representable.) When viewed under this duality, the state-passing translation becomes the
continuation-passing translation, as we now explain.
In Figure 5, we provide an internal language for enriched categories with powers. We
call this the ‘CPS variant of ECBV’, because it is the target of the continuation passing
translation (see below). The key difference with Figure 1 is that we have replaced the tensor
type (!A ⊗ B) by a power type (A → B). It is another fragment of the enriched effect calculus.
If V is a category with products and C is a category enriched in V with powers, then we can
interpret this CPS variant of ECBV in (V, C) through a variation of the interpretation in
def
Section 2.2. The power type is interpreted using the categorical powers: [[A → B]] = [[B]][[A]] .
A computation judgement Γ | ∆ ` t : A is interpreted as a morphism [[t]] : [[∆]] → [[A]][[Γ]]
in C.
Following the categorical analysis above, we define a bijection (−)◦ between the types
of ECBV with this CPS variant:
def

def

α◦ = α

def

1◦ = 1

(A × B)◦ = (A◦ × B◦ )

(A ( B)◦ = (B◦ ( A◦ )

(!A ⊗ B)◦ = A◦ → B◦

def

def

(7.1)

LINEAR USAGE OF STATE

25

This bijection extends to terms-in-context straightforwardly, and the resulting translation is
a restriction of the linear-use cps involution of the enriched effect calculus studied in [9, 10].
We achieve a linear-use continuation-passing translation by composing the state-passing
translation of Section 4 with this bijection (7.1). For clarity, we now write this out explicitly.
We fix a computation type R of ECBV, thought of as a return type. We translate FGCBV
types σ to ECBV value types σ R :
def

def

αR = α (σ × τ )R = σ R × τ R

def

def

1R = 1 (σ * τ )R = ((τ R ) → R) ( ((σ R ) → R)

We extend this translation to type contexts, taking an FGCBV type context Γ to an ECBV
type context ΓR .
The translation on terms is syntax-directed. We pick a variable k, completely fresh.
The translation takes an FGCBV value type judgement Γ `v V : σ to an ECBV judgement
ΓR | − ` V R : σ R , and it take an FGCBV producer judgement Γ `p M : σ to an ECBV
R
judgement ΓR | k : σ R → R ` Mk : R.
def

xR = x

def

?R = ?

def

hV, W iR = hV R, W R i
R def

R def

(return V )k = k (V R )
def

def

def

(π1 (V ))R = π1 (V R )

(π2 (V ))R = π2 (V R )
R

R

(M to x. N )k = (λk. Mk ) [λx. Nk ]
R

(λx. N )R = λk. λx. Nk

R def

(V W )k = (V R [k]) W R

The continuation-passing translation inherits the following results from the soundness
and full completeness of the state-passing translation (Theorems 4.1 and 4.3).
Proposition 7.1. For any computation type R, the continuation-passing translation is
sound:
(1) If Γ `v V ≡ W : σ then ΓR | − ` V R ≡ W R : σ R .
R
R
(2) If Γ `p M ≡ N : σ then ΓR | k : σ R → R ` Mk ≡ Nk : R.
Proposition 7.2 ([9], Corollary 1). Let R be a computation type constant. The continuationpassing translation is fully complete, in the following sense.
(1) If Γ `v V, W : σ and ΓR | − ` V R ≡ W R : σ R then Γ `v V ≡ W : σ.
R
R
(2) If Γ `p M, N : σ and ΓR | k : σ R → R ` Mk ≡ Nk : R then Γ `p M ≡ N : σ.
(3) For any ΓR | − ` t : σ R there is Γ `v V : σ such that ΓR | − ` t ≡ V R : σ R .
(4) For any ΓR | k : σ R → R ` t : R there is a producer term Γ `p M : σ such that
ΓR | k : σ R → R ` t ≡ M R : R.
The full completeness result is the same as the one obtained by Egger et al. [9, Corollary 1]
except that loc. cit. describes a translation on the full enriched effect calculus rather than
this fragment of it.
7.1. Sums and products. The CPS variant of the enriched call-by-value calculus can be
extended so that value types are closed under sums and computation types are closed under
products. Thus the types are as follows:
A, B ::= α | 1 | A × B | A ( B | 0 | A + B
A, B ::= α | A → B | 1 | A × B .

26

R. E. MØGELBERG AND S. STATON

(For brevity, we omit the term language, which is a fragment of the enriched effect calculus.)
The type system is designed to be dual to the enriched call-by-value language with sums.
The translation from that language to this one (7.1) is extended as follows:
def

(0)◦ = 0

def

(A + B)◦ = A◦ + B◦

def

(0)◦ = 1

def

(A ⊕ B)◦ = A◦ × B◦

and the analysis of Section 6.3 can be converted to a fully-complete linear-use continuationpassing style translation from fine-grain call-by-value with sums.
8. Effect theories
To illustrate the nature of the state-passing translation we endow our calculi with effects.
We do this is in a general way, by following the programme of Plotkin, Power and others [38]
whereby a theory of effects is presented as an algebraic theory.
We discuss how to add effects to the source and target languages of the state-passing
translation, FGCBV and ECBV. Our central observation is that to accommodate an
algebraic theory of effects in the enriched call-by-value calculus it is necessary and sufficient
to supply the chosen state type S with the structure of a comodel. The idea of state being
a comodel, particularly for the theory of store, arose in the work of Plotkin, Power and
Shkaravska [40, 35].
The section is structured as follows. We begin with an overview in which we study the
situation for a particular algebraic effect. We then proceed to look at a general notion of
effect theory (§8.3), its relationship with the state-passing translation (§8.4) and notions of
model and comodel (§8.5,8.6).
8.1. Overview. We give an overview of the situation in terms of a particular algebraic
theory: an algebraic theory for accessing a single bit of memory. This is an algebraic theory
in the classical sense (like monoids, groups, rings, etc.). It has one binary operation (?), a
unary operation (f) and the following four equations:
(v ? x) ? (y ? z) ≡ v ? z

x≡x?x

f(f(x)) ≡ x

f(x ? y) ≡ f(y) ? f(x)

(8.1)

Here is some intuition. If x and y are computations, then x ? y is the computation that first
reads the bit in memory and then branches to x or to y depending on whether the bit was
set. If x is a computation then f(x) is the computation that first flips the bit in memory (0
def
to 1, 1 to 0) and then continues as x. There are derived operations update0 (x) = x ? f(x)
def
and update1 (x) = f(x) ? x, which first write 0 (resp. 1) and then continue as x.
We now explain how to accommodate this algebraic theory in the fine-grain and enriched
call-by-value calculi.
8.1.1. Fine-grain call-by-value: algebraic operations and generic effects. In the fine-grain
call-by-value calculus (§3), the algebraic theory (8.1) can be accommodated in two equivalent
ways: by adding algebraic operations and by adding generic effects. In adding the operations,
we add the following term formation rules for each type σ:
Γ `p M : σ
Γ `p M : σ Γ `p N : σ
(8.2)
Γ `p fσ (M ) : σ
Γ `p M ?σ N : σ

LINEAR USAGE OF STATE

27

We also add the equations in (8.1) at each type, and an algebraicity equation for each
operation (e.g. [43, Def 3.14]):
Γ `p M1 : σ Γ `p M2 : σ Γ, x : σ `p N : τ
Γ `p (M1 ?σ M2 ) to x. N ≡ (M1 to x. N ) ?τ (M2 to x. N ) : τ
Γ `p M : σ Γ, x : σ `p N : τ
Γ `p fσ (M ) to x. N ≡ fτ (M to x. N ) : τ
The result is a programming language with higher types and a single bit of storage.
The second way to accommodate the algebraic theory into the fine-grain call-by-value
calculus is by adding generic effects. For this we need sum types (§6.2). The idea is that an
expression in n variables in the algebraic theory corresponds to a ground producer term of
type n (= 1 + · · · + 1). Thus we add the following axioms for term formation:
Γ `p deref(?) : 2

Γ `p flip(?) : 1

(8.3)

Informally, deref(?) is a computation that returns the boolean value of the memory cell, and
flip(?) is a computation that flips the value in the memory cell. An important observation
of Plotkin and Power [38] is that the algebraic operations can be recovered at all types from
the generic effects, as follows:
def

def

M ?σ N = ifp deref(?) then M else N

fσ (M ) = flip(?); M

where we use some shorthand notation:
def

ifp deref(?) then M else N = casep deref(?) of (in1 (x1 ).M |in2 (x2 ).N )
def

flip(?); M = flip(?) to x. M
Conversely the generic effects can be derived from the algebraic operations:
def

deref(?) = inp1 (?) ?2 inp2 (?)

def

flip(?) = f1 (?)

(The subscript 1 on f1 (?) is the unit type.) We can thus write the four equations (8.1)
directly in terms of generic effects:
`p deref(?) to x. deref(?) to y. return hx, yi ≡ deref(?) to x. return hx, xi : 2 × 2
`p return (?) ≡ deref(?); return (?) : 1

`p flip(?); flip(?) ≡ return (?) : 1

`p flip(?); deref(?) ≡ deref(?) to x. flip(?); return (¬x) : 2
(8.4)
writing ¬x for if x then in2 (?) else in1 (?).
The two derived operations for writing a bit can be combined into a single command assign:
Γ `v V : 2
def

Γ `p assign(V ) = ifp (deref(?) to x. return (V xor x)) then flip(?) else return (?) : 1
where xor is the evident binary operation on values of type 2. Using this derived command,
the four equations for accessing the bit of memory can be equivalently written as the three

28

R. E. MØGELBERG AND S. STATON

program equations of Plotkin and Power [37]:
− `p return (?) ≡ deref(?) to x. assign(x) : 1
p

x : 2 ` assign(x); deref(?) ≡ assign(x); return (x) : 2
x, y : 2 `p assign(x); assign(y) ≡ assign(y) : 1

(8.5)
(8.6)
(8.7)

which state that reading a cell and then writing the same value is the same as doing nothing,
that writing and the reading yields the value just written, and that the effect of two writes
equals that of the second.
8.1.2. Enriched call-by-value and state access operations. How can we accommodate the
algebraic theory for a bit of memory (8.1) in the enriched call-by-value calculus? In this
section we develop the following observation. Whereas in FGCBV each type should be a
model of the theory, in that (8.2) gives terms (?) and (f) at each type σ, in ECBV the
distinguished state type S should be a comodel of the theory, which means that there are
ground value terms
read : S ( S ⊕ S
and
flip : S ( S
(8.8)
which we call state access operations. (It is called a comodel because the arrows have
been reversed and (×) has become (⊕).) Using the isomorphism (S ⊕ S) ∼
= (!2 ⊗ S), we can
understand the type of (read) as S ( !2 ⊗ S. The idea is that the read operation takes a state
and returns the value stored in that state. It also returns a state: this is necessary because
state is linear and cannot be discarded or duplicated. Notice that, under the state-passing
translation, the two generic effects (8.3) become the two state access operations.
The four equations (8.1) are also appealing when written in terms of state access
operations.
− | s : S ` read[s] to (!b ⊗ s0 ). read[s0 ] to (!b0 ⊗ s00 ). !hb, b0 i ⊗ s00
≡ read[s] to (!b ⊗ s0 ). !hb, bi ⊗ s0 : !(2 × 2) ⊗ S
− | s : S ` s ≡ read[s] to (!b ⊗ s0 ). s0 : S

(8.9)

− | s : S ` flip[flip[s]] ≡ s : S
− | s : S ` read[flip[s]] ≡ read[s] to (!b ⊗ s0 ). !¬b ⊗ flip[s0 ] : !2 ⊗ S
Notice that the second equation says that the read operation does not change the state.
The two derived operations for writing a bit can be combined into a single state access
operation:

def
write = λx.xto(!b ⊗ s). read[s]to(!b0 ⊗ s0 ). (if (b xor b0 ) then (λs.flip[s]) else (λs.s))[s0 ]
: !2 ⊗ S ( S
Intuitively, write[!b ⊗ s] writes the value b to the state s, returning the updated state.
In Section 4 we have seen a fully-complete state-passing translation from FGCBV to
ECBV. This translation extends to a fully-complete translation from FGCBV with generic
effects to ECBV with state access operations.

LINEAR USAGE OF STATE

29

8.1.3. Continuation passing style and algebraic operations. Finally, we turn to the linear-use
continuation-passing style. In this setting, it is natural to require that the distinguished
return type R be a model of the theory. This is dual to the situation with state-passing
style, where the distinguished state type S is a comodel of the theory.
More precisely, we extend the CPS variant of ECBV (§ 7) with the theory of memory
access by adding ground value terms
(?) : R × R ( R

and

(f) : R ( R

satisfying the following equations:
− | k : (R × R) ×(R × R) ` ((π1 (π1 k)) ? (π1 (π2 k))) ? ((π2 (π1 k)) ? (π2 (π2 k)))
≡ (π1 (π1 k)) ? (π2 (π2 k)) : R
− |k: R ` k ≡ k ? k: R
− | k : R ` f[f[k]] ≡ k : R
− | k : R × R ` f[(π1 k) ? (π2 k)] ≡ (f[π2 k]) ? (f[π1 k]) : R
Thus the generic effects in the source language endow the return type R of the linear-use
continuation-passing translation with the structure of a model for the algebraic theory.
8.1.4. Further simple examples of algebraic theories for computational effects. The theory
of accessing a bit of memory is perhaps the simplest example of a stateful effect. The
connections between algebraic operations, generic effects and state access operations also
work for less state-like effects.
Printing. The algebraic theory of printing a single bit has two unary function symbols, p0
and p1 . For instance, the term p0 (p1 (x)) should be understood as the computation that
first prints 0, then prints 1, then continues as x. There are no equations in this theory.
The generic effects for printing can be grouped together into one producer term
x : 2 `p print(x) : 1
thought of as a command that prints its argument.
As a state access operation, we have a function print : !2 ⊗ S ( S which, given a bit
and a state, returns a new state. Intuitively, S is a list of everything printed so far, and
print appends its first argument to its second argument.
Probability. There are different algebraic theories of probabilistic choice. The simplest one
is the theory of ‘mean-values’ considered by Heckmann [16] (but perhaps first introduced
by Aczél [2]): it has one binary function symbol
and its axioms are the medial law,
idempotence and commutativity:
(u

x)

(y

z) ≡ (u

y)

(x

z)

x≡x

x

x

y≡y

x

The idea is that a computation x y tosses a coin, and proceeds as x if heads and y if tails.
The generic effect for is Γ `p toss(?) : 2 which, intuitively, tosses the coin and returns
the result. In this style, the first equation is written
− `p toss(?) to x. toss(?) to y. (x, y) ≡ toss(?) to y. toss(?) to x. (x, y) : 2 × 2
It says that it doesn’t matter which order you toss coins.

30

R. E. MØGELBERG AND S. STATON

The state access operation toss : S ( !2 ⊗ S can be thought of as making the coin an
explicit parameter: we can think of S as a type of coins. In this style, the second equation
− | s : S ` s ≡ toss[s] to (!b ⊗ s0 ). s0 : S
says that when you toss a coin you get the same coin back. The third equation
− | s : S ` toss[s] ≡ toss[s] to (!b ⊗ s0 ). (!(¬b) ⊗ s0 ) : !2 ⊗ S
says that if you toss a coin it is the same as tossing a coin and turning it once more without
looking. This illustrates that probability is not really stateful and so for this effect the
approach based on algebraic operations is perhaps the most profitable perspective. The point
is that different computational effects are better suited to different approaches (algebraic
operations, generic effects, and state access operations) even though all three approaches
are always available.
8.2. State access operations, algebraic operations, and generic effects. We now
make precise the informal connections made between state access operations, generic effects
and algebraic operations in the previous section. We do this by revisiting the results of
Plotkin and Power [38] in the context of an enriched model (V, C).
In the previous section we focused on the classical situation where arities are natural
numbers. However, from the perspective of state access operations, generic effects and
algebraic operations have little to do with natural numbers per se. It is just as easy to allow
arities to be arbitrary objects of the base category V. In the following result we do not make
an artificial restriction to natural numbers. Instead we consider an operation with arity
[A1 , . . . , An ] and a parameter from B, where A1 , . . . , An , B are objects of V. The classical
case of an n-ary operation is recovered by setting the objects A1 , . . . , An , B to all be the
terminal object 1.
Theorem 8.1. Let (V, C) be an enriched model with sums. Let S be an object of C. Let
A1 . . . An and B be objects of V. The following data are equivalent:
(1) A state access operation: a morphism B · S → A1 · S ⊕ . . . ⊕ An · S in C.
(2) A generic effect: a morphism B → TS (A1 + · · · + An ) in V, where TS is the monad
C(S, (−) · S).
(3) An algebraic operation: a V-natural family of morphisms in V
 Qn
Ai
→ (US X)B X∈C
i=1 (US X)
def

where US (X) = C(S, X).
The last point requires some explanation. First, even though V is not cartesian closed,
exponentials with base US (X) exist: (US X)A ∼
= C(A · S, X). Second, the constructions
Q
def
def
n
Ai
F =
G = (US (−))B
i=1 (US (−))
can be understood as V-functors F, G : C → V, since there are families of morphisms
{FX,Y : C(X, Y ) × F (X) → F (Y )}X,Y ∈C

{GX,Y : C(X, Y ) × G(X) → G(Y )}X,Y ∈C

in V that satisfy the laws for functors (respecting identities and composition). Thirdly, a
family of morphisms {φX : F (X) → G(X)}X∈C is called V-natural if the following diagram

LINEAR USAGE OF STATE

31

commutes in V for all X and Y :
C(X, Y ) × F (X)
FX,Y

C(X,Y )×φX

/ C(X, Y ) × G(X)



F (Y )



φY

GX,Y

/ G(Y )

(It is perhaps more compelling if algebraic operations are defined as structure on a V-category
of TS -algebras, but this V-category cannot be constructed without further assumptions on
V — see [38, §7].)
Proof of Theorem 8.1. To see the connection between (1) and (2), consider the following
bijections:
HomC (B · S, A1 · S ⊕ . . . ⊕ An · S) ∼
= HomV (B, C(S, A1 · S ⊕ . . . ⊕ An · S))
∼
= HomV (B, C(S, (A1 + · · · + An ) · S))
= HomV (B, TS (A1 + · · · + An )).
To see the connection between (1) and (3), we note that
Qn
Ai ∼
and
= C(A1 · S ⊕ · · · ⊕ An · S, X)
i=1 (US X)

(US X)B ∼
= C(B · S, X)

and the enriched Yoneda lemma gives
V-Nat(C(A1 · S ⊕ · · · ⊕ An · S, −), C(B · S, −)) ∼
= HomC (B · S, A1 · S ⊕ · · · ⊕ An · S).
We remark that the equivalence of (2) and (3) is essentially Theorem 2 of [38].
8.3. Effect theories. In the previous section we described connection between state access
operations, generic effects and algebraic operations. As we explained, the natural level of
generality for this is more sophisticated than the classical setting: the arity of an operation
is a list and we allow the operation to take a parameter. This suggests a generalization of
algebraic theories that we call ‘effect theories’, since they are useful from the computational
perspective.
The illustration in Section 8.1 involves storage of a single bit. A motivating example of
effect theory arises from modifying that theory above to allow storage of a more interesting
datatype. In FGCBV, we would like to have an (abstract) type Val of storable values, and
generic effects deref and assign with typing judgements
Γ `v V : Val
(8.10)
Γ `p deref(?) : Val
Γ `p assign(V ) : 1
We add to the theory of equality for FGCBV (Fig. 2) the three equations for global store
proposed by Plotkin and Power [37] (8.5–(8.7)):
`p return (?) ≡ deref(?) to x. assign(x) : 1
x : Val `p assign(x); deref(?) ≡ assign(x); return (x) : Val
x, y : Val `p assign(x); assign(y) ≡ assign(y) : 1
Our notion of effect theory accommodates the classical kinds of theory in the overview
(§ 8.1) and also the more general kind of theory of memory access illustrated above. It
is roughly the same as that used by Plotkin and Pretnar [36, §3]. The main difference is

32

R. E. MØGELBERG AND S. STATON

in the presentation: we use generic effects rather than algebraic operations. Rather than
introducing a new calculus for expressing the allowable equations of an effect theory, we use
the first-order fragment of FGCBV.
Value theories. Before we introduce effect theories we briefly discuss value theories, which
are simple extensions of the value judgements of FGCBV. By a value signature we shall
simply mean a signature for a many-sorted algebraic theory in the usual sense. This means
a set of type constants ranged over by α, β, and a set of term constants f with a given arity
f : (α1 , . . . , αn ) → β, where the αi , β range over type constants. We can extend FGCBV
along a value signature by adding the type constants and the typing rule
Γ `v ti : αi (i = 1, . . . , n)
(8.11)
Γ `v f (t1 , . . . , tn ) : β
for every term constant f : (α1 , . . . , αn ) → β in the signature. A value theory is a value signature with a set of equations, i.e. pairs of terms typable in the same context Γ `v V = W : β,
where V, W are formed only using variable introduction and the rule (8.11).
Effect theories. An effect signature consists of a value theory and a set of effect constants
each with an assigned arity e : β̄; ᾱ1 + . . . + ᾱn consisting of a list β̄ of type constants and
a formal sum of lists of type constants, ᾱ1 + . . . + ᾱn . (Here we are abbreviating a list
(β1 . . . βm ) using the notation β̄, etc.) FGCBV can be extended along an effect signature by
adding, for every effect constant e : β̄; ᾱ1 + . . . + ᾱn , a typing judgement
Γ `v V1 : β1 . . . Γ `v Vm : βm
(8.12)
Γ `p e(V1 , . . . , Vm ) : ᾱ1 + . . . + ᾱn
where β̄ = (β1 , . . . , βm ). In the conclusion, the vectors ᾱi should be understood as the
product of the types in the vector.
Here are some examples of effect signatures:
• The theory of reading/flipping a bit of memory (§ 8.1) has no value type constants. It has
two effect constants, deref and flip. The effect constant deref has arity 1; 1 + 1 and
the effect constant flip has arity 1; 1, where 1 is the empty string.
• The theory for storing an abstract datatype (8.10) has one value type constant Val and a
pair of effect constants (deref : 1; Val) and (assign : Val; 1). In this case term constants
in the value theory can be used to add basic operations manipulating values in Val: we
could ask that the storable values form a ring. (In future, it would be interesting to allow
Val to have more structure, for example as an inductive type such as the natural numbers,
but it is not clear how to extend the proof of Theorem 10.2 to inductive types.)
An effect theory comprises an effect signature and a set of equations. The equations are
pairs of producer terms-in-context Γ `p M ≡ N : τ of a restricted kind: they must be built
from the first-order fragment of fine-grain call-by-value in Figure 6. This notion of ‘effect
theory’ is connected with the classical notion of algebraic theory in Section 8.1 as follows. If
the value theory is empty, with no type constants (α) and no function symbols, then the lists
of type constants in (8.12) must all empty, and each generic effect is an operation of arity n.
This is the generic effect presentation of an algebraic theory as described in Section 8.1.1.

LINEAR USAGE OF STATE

33

Types.
σ, τ ::= α | 1 | σ × τ | 0 | σ + τ

Terms. The grammar for terms is as follows. Typing judgements are in Figure 2, Figure 4,
and Equations (8.11) and (8.12).
V ::= x | f (V1 , . . . , Vn ) | ? | π1 (V ) | π2 (V ) | hV1 , V2 i
| ?(V ) | in1 (V ) | in2 (V ) | case V of (in1 (x1 ).W1 |in2 (x2 ).W2 )
M ::= return V | M to x. N | casep M of (in1 (x1 ).N1 |in2 (x2 ).N2 )
Figure 6: The sub-calculus of fine-grain call-by-value that is used for describing effect theories
8.4. Effect theories and the state-passing translation. An effect theory is an extension
to the fine-grain call-by-value language. In Section 4 we explained how the linear-use statepassing translation goes from FGCBV to ECBV. We now explain how ECBV needs to
be extended to support this state-passing translation. The restricted nature of the effect
theories makes this particularly straightforward and illuminating.
Value theory:
• For each type constant in the value theory, we assume a value type constant.
• For each term constant f : ᾱ → β in the value theory we add the following term formation
rule:
Γ | − ` t1 : α1
...
Γ | − ` tn : αn
Γ | − ` f (t1 , . . . , tn ) : β
• An equation in the value theory must only be formed from variable introduction and the
rule (8.11). Thus each equation Γ `v V ≡ W : β in the value theory can be understood as
an equation Γ | − ` V ≡ W : β between value judgements.
Effect signature:
• We assume a chosen computation type constant S.
• For each effect constant e : β̄; ᾱ1 + . . . + ᾱn we add a constant at value type,
− | − ` e : !β̄ ⊗ S ( !(ᾱ1 + . . . + ᾱn ) ⊗ S

(8.13)

For the theory of reading/flipping a bit, this yields the constants read and flip in (8.8).
For the theory of storing an abstract datatype, this yields two state access operations:
read : S ( !Val ⊗ S

write : !Val ⊗ S ( S

State-passing translation:
• Recall that the state-passing translation (§4) takes a producer judgement of FGCBV
S
Γ `p M : σ to a computation judgement of ECBV ΓS | s : S ` Ms : !σ S ⊗ S. We extend
the state-passing translation to operate on effects:
def

S

(e(V1 . . . Vm ))Ss = e(!(V1 , . . . , VmS ) ⊗ s)

34

R. E. MØGELBERG AND S. STATON

• We use this extended state-passing translation to translate the equations in the effect
theory into ECBV, in such a way that the extended translation is sound by construction.
Each equation in the effect theory Γ `p M ≡ N : τ becomes an equation in ECBV:
Γ | s : S ` MsS ≡ NsS : !τ ⊗ S.
Notice that we do not need to translate the types in Γ because equations in an effect
theory must be from the first-order fragment of fine-grain call-by-value (Fig. 6) which is
shared with enriched call-by-value. For instance, the equations in the effect theory for
reading/flipping a bit (8.4) give rise to the equations on the state object (8.9). The three
equations for storing an abstract datatype (8.5–8.7) become the following equations for a
state object S:
− | s : S ` s ≡ write[read[s]] : S
x : Val | s : S ` read[write[!x ⊗ s]] ≡ !x ⊗ (write[!x ⊗ s]) : !Val ⊗ S
x, y : Val | s : S ` write[!y ⊗ write[!x ⊗ s]] ≡ write[!y ⊗ s] : S
8.5. Models and comodels of effect theories. Our analysis of effect theories in Sections 8.3 and 8.4 has been syntactic. We now provide a model-theoretic treatment. We define
the interpretation of effect theories in Kleisli models (§8.5.2) and enriched models (§8.5.3).
We then define what it means to be a model of an effect theory in general terms (§8.5.4).
8.5.1. Models of value theories. Let V be a distributive category. An interpretation of a
value signature in V is given by interpretations of the type constants α as objects [[α]]
of V, and interpretations of term constants f : ᾱ → β as morphisms [[f ]] : [[ᾱ]] → [[β]]. (Here,
def
if ᾱ = (α1 , . . . , αn ) then [[ᾱ]] = [[α1 ]] × · · · × [[αn ]].) This interpretation is extended to
interpret a term in context Γ `v V : β as a morphism [[V ]] : [[Γ]] → [[β]]. An interpretation of
a value theory is an interpretation of the signature such that [[V ]] = [[W ]] for each equation
Γ `v V ≡ W : β in the value theory.
8.5.2. Interpreting effect theories in Kleisli models. Let (V, C, J) be a distributive Kleisli
model (Def. 6.2) and suppose an interpretation of the type constants α, β is given. An
interpretation of an effect theory E in (V, C, J) is given by an interpretation of the value
theory in V and an interpretation of each effect constant e : β̄; ᾱ1 + · · · + ᾱn in E as a
morphism [[e]] : [[β̄]] → ([[ᾱ1 ]] + · · · + [[ᾱn ]]) in C, satisfying the equations of the theory.
8.5.3. Comodels of effect theories in enriched models. Let (V, C) be a distributive enriched
model in the sense of Definition 6.1. Thus V is a distributive category and C is a category
enriched in V with copowers and coproducts. A comodel of the effect theory in C is an
object S of C together with a morphism [[e]] : [[β̄]] · S → ([[ᾱ1 ]] + · · · + [[ᾱn ]]) · S in C for every
effect constant e : β̄; ᾱ1 + . . . + ᾱn such that for each equation Γ `p M ≡ N : τ in the effect
theory, the interpretations of M and N under the state passing style yield equal morphisms:
[[MsS ]] = [[NsS ]] : [[Γ]] · S → [[τ ]] · S.

LINEAR USAGE OF STATE

35

8.5.4. Models of effect theories in dual enriched models. We now justify our use of the term
‘comodel’ in Section 8.5.3 by showing that it is a generalization of the standard usage, i.e.,
dual to the concept of model of an algebraic theory known from classical algebra. Further
investigations of the notion of effect theory used in this paper along with relations to existing
notions of enriched algebraic theories [23, 39] could be an interesting topic of further research.
To dualize the notion of comodel, let (V, Cop ) be a distributive enriched model, i.e.,
let V be a distributive category and C a category enriched in V with powers and products, as
in Section 7. A model of the effect theory in C is a comodel in Cop . Explicitly this amounts
to an object R of C together with a morphism [[e]] : R([[ᾱ1 ]]+···+[[ᾱn ]]) → R[[β̄]] between powers
in C for every effect constant e : β̄; ᾱ1 + . . . + ᾱn such that for each equation Γ `p M ≡ N : τ
in the effect theory, the interpretations of M and N in continuation passing style yield
equivalent morphisms:
R
R
[[Mk ]], [[Nk ]] : R[[τ ]] → R[[Γ]] .
Because the terms of the effect theory are of a restricted kind, it is straightforward
to directly describe the interpretations of effect terms as morphisms between powers of
R, by induction on the structure of typing derivations. For instance, consider the casep
R
R
rule in (6.1). Given interpretations JMk K : RJσ1 +σ2 K → RJΓK and J(Ni )k K : RJτ K → RJΓ,σi K
R
(i = 1, 2), the interpretation J(casep M of (in1 (x1 ).N1 |in2 (x2 ).N2 ))k K is the composite
R

R

R JΓK

∆

(J(N1 ) K,J(N2 ) K)
JMk K
R
RJτ K −−−−−k−−−−−k−→ RJΓ,σ1 K × RJΓ,σ2 K ∼
= R(Jσ1 +σ2 K)×JΓK −−−−−→ RJΓK×JΓK −−→ RJΓK .
R

def

As another example, J(return V )k K = RJV K : RJτ K → RJΓK if Γ `v V : τ .
We now return to the setting of classical algebra, when the value theory has no type
constants or value constants. We will show that models in the above sense are models of
algebraic theories in the classical sense. If there are no type constants then every type in
the effect language (Figure 6) is isomorphic to one of the form 1 + 1 + · · · + 1. The arity
of an effect constant e : β̄; ᾱ1 + . . . + ᾱn must comprise β̄ as the empty list (since there
are no type constants) and ᾱ1 + . . . + ᾱn must be a sequence of n empty lists. Thus the
interpretation of e in the model is a morphism [[e]] : Rn → R.
We now explain the nature of equations in this restricted setting. In what follows,
we will write a natural number n for the type that is the n-fold sum of 1. We will write
ini (?) for the ith injection of type n, where 1 ≤ i ≤ n, and we will make use of n-ary case
constructions, case V of (in1 (?).W1 | . . . |inn (?).Wn ) to destruct terms V of type n. These
are just syntactic shorthand for terms that can be defined in the language in Figure 6.
We can focus on equations where Γ is empty. This is because every context has a finite
number of ground valuations — if a variable x in a context has type n, then it could be
valued with in1 (?) . . . inn (?) — and, moreover, an effect equation Γ `p M ≡ N : n is satisfied
if and only if it is satisfied at each ground instantiation.
The next step is to note that every effect term − `p M : 1 + 1 · · · + 1 is equal to one
built from the following rules:
− `pn M1 : n . . . − `pn Mm : n
(e : −; m)
− `pn casep e(?) of (in1 (?).M1 | . . . |inm (?).Mm ) : n
− `pn inpi (return ?) : n

1≤i≤n

36

R. E. MØGELBERG AND S. STATON

It is informative to look at the interpretation of these normalized terms. Given interpretations
[[M1 ]], . . . , [[Mm ]] : Rn → R, we have
([[M1 ]],...,[[Mm ]])

[[e]]

[[casep e(?) of (in1 (?).M1 | . . . |inm (?).Mm )]] = Rn −−−−−−−−−−→ Rm −−→ R,
π

i
[[inpi (return ?)]] = Rn −→
R.

Thus we see that, in the situation where there are no type constants or value constants, the
new general notion of model is the classical notion of model for an algebraic theory.
8.6. Examples of set-theoretic models and comodels. We revisit the simple example
effect theories from Sections 8.1 and 8.3 from the model-theoretic perspective. In each case,
we find that there are comodels that are state-like.
8.6.1. Storage. The category Set is enriched in itself with copowers given by products and
the enrichment given by the function space. The set 2 = {0, 1} is a comodel for the theory of
accessing a bit of memory (§8.1), with read(x) = (x, x) and flip(x) = ¬x. This is a comodel
for the theory in the enriched model (Set, Set). Power and Shkaravska [40] showed that 2
is the final comodel in Set for the theory of accessing a bit of memory.
As an aside, we note that Set is actually equivalent to the category of models for the
theory of accessing a bit of memory. The theory of reading a bit of memory is sometimes
called the theory of ‘rectangular bands’ because every model is isomorphic to one of the
def
form X × Y , with (x, y) ? (x0 , y 0 ) = (x, y 0 ). The anti-involution operation (f) enforces
that the model is isomorphic to one of the form X × X, and thus determined by a single
set. This phenomenon has been investigated in a more general setting by Métayer [30] and
Mesablishvili [29].
We can consider set theoretic models of the theory of storing an abstract datatype (8.10).
What is needed is an interpretation Val of the value sort, which also plays the role of the
state object. We let read(x) = (x, x) and write(v, s) = v. This is a comodel for the theory
for store in the enriched model (Set, Set).
In both cases, the induced monad on Set is the store monad ((−) × S)S .
8.6.2. Printing. Let 2∗ -Act be the category of algebras for the theory of printing a bit. The
objects are triples (X, p0 , p1 ) where p0 , p1 : X → X, and the morphisms (X, pX,0 , pX,1 ) →
(Y, pY,0 , pY,1 ) are functions X → Y that commute with the operations. As any ordinary
category, this category is enriched in Set. It has copowers given by
def

A · (X, pX,0 , pX,1 ) = (A × X, p(A·X),0 , p(A·X),1 )

def

where p(A·X),i (a, x) = (a, pX,i (x)).

Thus (Set, 2∗ -Act) is an enriched model.
The algebra structure of each algebra equips it with the structure of a comodel in
the category of algebras. The leading example is the set 2∗ of strings over {0, 1}, with
p2∗ ,i (s) = si. The induced state monad 2∗ -Act(2∗ , (−) · 2∗ ) is isomorphic to the monad
2∗ × (−) on Set. We can understand a string in 2∗ as a state: it is the list of things output
so far.

LINEAR USAGE OF STATE

37

8.6.3. Probability. Let MVAlg be the category of mean-value algebras. The objects are pairs
(X, ) of a set X and a binary operation satisfying the laws of mean-value algebras (§8.1.4).
The pair (Set, MVAlg) is an enriched model.
The one-element set is trivially a mean-value algebra, and it can be given the structure
of a comodel in the category of mean-value algebras. We can understand the one-element set
as a set of states: this captures the idea that probability is a stateless notion of computation.
Nonetheless, this ‘state object’ induces a ‘state monad’ on Set. This can be understood as a
monad D of finite dyadic probability distributions. By a finite dyadic probability distribution
on a set X,
P we mean a function p : X → [0, 1] such that supp(p) = {x ∈ X | p(x) 6= 0}
is finite, x∈supp(p) p(x) = 1, and for all x, p(x) has a finite binary representation. The
monad D has D(X) as the set of all finite dyadic probability distributions; the unit
picks out the Kronecker distributions, and multiplication µX : D(D(X)) → D(X) takes a
distribution p : D(X) → [0, 1] on D(X) to a distribution µX (p) : X → [0, 1] on X, given by
def P
µX (p)(x) =
q∈supp(p) (p(q) × q(x)).
In general, when the state object is a terminal object then the induced monad preserves terminal objects. A terminal-object-preserving monad is sometimes called affine [25,
Thm. 2.1] and the corresponding effects are said to be discardable (e.g. [12], [47, Def. 4.2.4])
since the following rule is admissible in the fine-grain call-by-value language.
Γ `p t : A Γ `p u : B
(x not free in u)
Γ `p t to x. u ≡ u : B
8.7. Relating notions of (co)model for effect theories. We now extend Theorem 8.1
to show, for each effect theory E, a bijective correspondence between the following: comodel
structures on S, interpretations of E in the Kleisli model (V, KlS , JS ), algebraic operations
equipping each US X with a model structure for E. The latter notion requires some
explanation, because the definition of model given in Section 8.5.4 defines only what it
means for R in a category C to be a model of E if (V, Cop ) is an enriched model, and in
the setting of Theorem 8.1 Vop is generally not V-enriched.
To generalize the notion of model, let V be a distributive category, let R be a fixed
object of V such that all exponents of the form RA exist (i.e. V(− × A, R) is representable
for all A) and let an interpretation of the value theory of E be given. We define a model
structure for E on R to be an interpretation of E in the Kleisli model (V, KlRR(−) , J) where
KlRR(−) has the same objects as V, but where a morphism A → B in KlRR(−) is a morphism
(−)

RB → RA in V. This is isomorphic to the Kleisli category for the strong monad RR on
V. By construction, the model structure interprets each effect constant e : β̄; ᾱ1 + . . . + ᾱn
as a morphism
[[e]] : R([[ᾱ1 ]]+···+[[ᾱn ]]) → R[[β̄]] .
If V is cartesian closed then (V, Vop ) is an enriched model and the above definition of a
model structure for E on R is equivalent to the one given in Section 8.5.4.
Theorem 8.2. Let (V, C) be an enriched model with sums, let S be an object of C, let E
be an effect theory and let an interpretation of the value theory of E in V be given. The
following data are equivalent:
(1) A comodel structure for E on S
(2) An interpretation of E in the Kleisli model (V, KlS , JS )

38

R. E. MØGELBERG AND S. STATON

(3) For each effect constant e : β̄; ᾱ1 + . . . + ᾱn in E an algebraic operation: a V-natural
family of morphisms in V
o
nQ
n
[[ᾱi ]]
[[β̄]]
→
(U
X)
X)
(U
S
S
i=1
X∈C

def

equipping each US (X) = C(S, X) with a model structure for E.
Proof. We first prove equivalence of (1) and (2). First note that in both cases, an effect
constant e : β̄; ᾱ1 + . . . + ᾱn is modelled as a morphism
[[e]] : [[β̄]] · S → ([[ᾱ1 ]] + · · · + [[ᾱn ]]) · S .
It thus suffices to show that for any term Γ `p M : σ of the fragment of Figure 6 the
S
S
S
morphisms [[Ms ]], [[M ]] : [[Γ]] · S → [[σ]] · S are equal, where [[Ms ]] is the ECBV term Ms
interpreted in the enriched model (V, C, S) and [[M ]] is the fine-grain call-by-value term
M interpreted in the Kleisli model (V, KlS , JS ). This can be proved by induction on the
structure of M .
To prove equivalence of (1) and (3) we build on the equivalence of state access operations
and algebraic operations of Theorem 8.1. First we show that for any term Γ `p M : σ of the
fragment of Figure 6 the equation
C([[MsS ]], X) = [[M ]] : C([[σ]] · S, X) → C([[Γ]] · S, X)

(8.14)

holds. This time the denotation brackets on the left hand side refers to the interpretation of
ECBV in (V, C, S), and the denotation brackets on the right hand side refer to the interpretation of fine-grain call-by-value in the Kleisli model (V, KlRR(−) , J), where R = US (X). As
S

above, [[Ms ]] is simply the interpretation of M in the Kleisli model (V, KlS , JS ). Equation
(8.14) can be proved by induction on the structure of M . (There is also a categorical
perspective on this, based around the construction C(−, X) which gives rise to an identityon-objects functor KlS → KlRR(−) that preserves sums and the action of V, although it
does not preserve the enrichment.)
From (8.14) we deduce the equivalence of (1) and (3). In fact (1) =⇒ (3) is immediate:
suppose S is a comodel, and consider an equation Γ ` M ≡ N : τ in the theory. Since S is a
S
S
comodel, we have [[Ms ]] = [[Ns ]] and so [[M ]] = [[N ]] as interpreted in (V, KlRR(−) , J).
For (3) =⇒ (1), suppose that R = C(S, X) is a model for every X, naturally in X.
Then by (8.14)
C([[MsS ]], X) = C([[NsS ]], X) : C([[τ ]] · S, X) → C([[Γ]] · S, X)
holds for all equations Γ ` M ≡ N : τ and all X. The enriched Yoneda embedding is full
S
S
and faithful and so [[Ms ]] = [[Ns ]] : [[Γ]] · S → [[τ ]] · S, proving that S is a comodel.
We remark that the equivalence of (2) and (3) is in the spirit of [38, §6].
8.8. Generalizing full completeness to the case of effects. The full completeness
result of Theorem 4.3 extends verbatim to the case of the calculi augmented with an effect
theory E. The proof is based on an extension of the coreflection theorem (Theorem 5.3)
which we state below.
First 2-categories dKleisliE and dEcbvE of distributive Kleisli models of E and distributive
enriched models of E are defined. These are defined similarly to Kleisli and Ecbv except

LINEAR USAGE OF STATE

39

morphisms are required to preserve coproducts (up to isomorphism) and the interpretation
of E (strictly). Details can be found in Appendix A.5.
def

def

Lemma 8.3. The assignments St(V, C, J) = (V, C, 1) and Kl(V, C, S) = (V, KlS , JS )
extend to 2-functors
St : dKleisliE → dEcbvE
Kl : dEcbvE → dKleisliE
Proof (sketch). We just show that these are well-defined on objects. The case of Kl is simply
the implication from (1) to (2) of Theorem 8.2.
In the case of St we must show that 1 carries a comodel structure for E whenever
(V, C, J) models E. An effect constant e : β̄; ᾱ1 + . . . + ᾱn can be modelled as the composite
J(π1 )

[[e]]

Jhid,! i

[[β̄]] · 1 −−−→ [[β̄]] −−→ [[ᾱ1 + . . . + ᾱn ]] −−−−→ [[ᾱ1 + . . . + ᾱn ]] · 1
where [[e]] refers to the interpretation of e in the given E-model structure of (V, C, J). We
must show that this defines a comodel, i.e., that the equations are satisfied. To this end
one can prove that for any term Γ `p M : σ of the fragment of fine-grain call-by-value used
S
for effect theories (Figure 6) the equation [[Ms ]] = J(hid , ! i) ◦ [[M ]] ◦ J(π1 ) holds. Here, on
the left hand side the double brackets refer to the interpretation of ECBV in (V, C, 1) and
the brackets on the right hand side refer to the interpretation of fine-grain call-by-value
in (V, C, J). This equation is proved by induction on typing derivations. Thus, for any
equation Γ `p M ≡ N : τ in E, since (V, C, J) models E we have [[M ]] = [[N ]], and thus also
S
S
[[Ms ]] = [[Ns ]], proving that 1 is indeed a comodel.
We end this section by stating the coreflection theorem for models of effect theories.
Theorem 8.4. The 2-functor St : dKleisliE → dEcbvE is left biadjoint to Kl, i.e., for any
pair of objects (V, C, J) and (V0 , C0 , S) of dKleisliE and dEcbvE respectively, there is an
equivalence of categories
dEcbvE (St(V, C, J), (V0 , C0 , S)) ' dKleisliE ((V, C, J), Kl(V0 , C0 , S))
natural in (V, C, J) and (V0 , C0 , S). Moreover, the unit of the adjunction η : id dEcbvE →
Kl ◦ St is an isomorphism.

9. Relationship with Atkey’s parameterized monads
Atkey’s work on parameterized monads [3], has proven relevant to functional programming (e.g. [24, §5.2]). In this section we show that parameterized monads are essentially the
same as enriched models.
Recall that, in general category theory, if a functor F : A × S → B is such that
F (−, S) : A → B has a right adjoint G(S, −) : B → A for each S, then these right adjoints
together form a functor G : S op × B → A called the parameterized right adjoint. Atkey
has carried out a study of a generalized form of monad that arises from parameterized
adjunctions: the functor G(−1 , F (−2 , −3 )) : S op × A × S → A is called a parameterized

40

R. E. MØGELBERG AND S. STATON

monad. Thus a parameterized monad is a functor T : S op × A × S → A together with
extranatural families of morphisms
ηS,A :A → T (S, A, S)
µS1 ,S2 ,S3 ,A :T (S1 , T (S2 , A, S3 ), S2 ) → T (S1 , A, S3 )
satisfying monad laws. A first example of a parameterized monad is the parameterized state
def
monad on the category of sets: T (S1 , A, S2 ) = [S1 ⇒ A × S2 ].
Every enriched model (V, C) contains a parameterized adjunction, since C(−1 , −2 ) :
Cop × C → V is by definition a parameterized right adjoint for (−1 ) · (−2 ) : V × C → C.
Conversely, in the theory of parameterized monads, the following Kleisli construction [3,
Prop. 1] plays a key role. Given a parameterized monad T : S op × A × S → A, the objects
of the Kleisli category are pairs (A, S) of an object of A and an object of S, and a morphism
(A, S) → (A0 , S 0 ) is a morphism A → T (S, A0 , S 0 ) in V. This is a first step towards building
an enriched model from a parameterized monad.
Plain parameterized monads are not especially relevant to the theory of programming
languages, just as plain monads are not very relevant. In his study of parameterized monads,
Atkey focuses on strong parameterized monads with Kleisli exponentials [3, §2.4.1]. He uses
these to provide semantics for a ‘command calculus’, which is a term language closely related
to our basic enriched call-by-value calculus (Figure 1).
Proposition 9.1. Let V be a category with finite products. The following data are equivalent.
(1) A strong parameterized monad on V with Kleisli exponentials, taking parameters in a
category S [3, §2.4.1].
(2) A category C enriched in V with copowers (i.e., an enriched model – §2.2) with a chosen
subcategory S of C such that every object X of C is of the form X = A · S for A in V
and S in S.
Proof notes. Given a strong parameterized monad T : S op × V × S → V, we let C be
the Kleisli category for T , as above. The pair (V, C) forms an enriched model with
def
A · (B, S) = (A × B, S): this is a rephrasing of what it means for T to be strong and have
Kleisli exponentials. Moreover S can be identified with a subcategory of C whose objects
are of the form (1, S) and whose morphisms are induced by the morphisms in S.
Conversely, suppose we are given an enriched model (V, C) and a chosen subcategory S
of C. We define a parameterized monad T : S op × V × S → V by
def

T (S, A, S 0 ) = C(S, A · S 0 )

.

It is routine to check that the two constructions are mutually inverse, up-to equivalence of
categories.

10. Relationship with the enriched effect calculus
Our enriched call-by-value calculus (ECBV) is a fragment of the enriched effect calculus (EEC,
[8, 7]) which was designed to analyze linear usage in effectful computation. We now show
that ECBV is not only a syntactic fragment of EEC: every model of the enriched call-by-value
calculus embeds in a model of the enriched effect calculus.

LINEAR USAGE OF STATE

41

The enriched effect calculus extends the enriched call-by-value calculus that we introduced
in Section 2 with some type constructions:
A, B ::= α | 1 | A × B | A ( B

| 0 | A + B | A → B | α | 0 | A ⊕ B | !A ⊗ B | !A

A, B ::= α | !A ⊗ B
|
{z

| 0 | A ⊕ B | 1 | A × B | A → B | !A .
|
| {z }
{z

Figure 1

}

Figure 3

Full EEC [8]

}

The additional types are: products (A × B) and powers (A → B) of computation types; an
operation to coerce a value type A into a computation type !A; a space of pure functions
(A → B) between value types; and an implicit inclusion of computation types as value types.
These additional types have been used to describe other aspects of effectful computation.
We briefly considered a linear-use CPS translation in Section 7, based on [9, 10], for which
we needed products and powers of computation types. Egger et al. [8] also describe monadic
call-by-name and call-by-value interpretations, for which they use the coercion of value types
into computation types and the implicit inclusion of computation types in value types.
The additional types of EEC do not affect the full completeness of the linear state-passing
translation (Thm. 4.3), for the following reason. In Theorem 10.2 we show that every model
of ECBV embeds in a model of EEC; conservativity of EEC over ECBV then follows from
a strong normalisation result for EEC [7]. Thus the linear-use state-passing translation of
Section 4 can be understood as a fully complete translation into EEC.
Definition 10.1. A closed enriched model is a pair of categories (V, C) such that V is
cartesian closed with coproducts and C is V-enriched with powers and copowers and products
and coproducts.
A model of EEC (V, C, F, U ) [8] is a closed enriched model (V, C) together with a
V-enriched adjunction F a U : C → V.
We refer to [8] for the term calculus and interpretation of EEC in EEC models. Here,
we will analyze how the notion of EEC model compares to the other notions of model that
we have discussed so far. One simple observation is that every closed enriched model is a
distributive enriched model in the sense of Definition 6.1. Another observation is that the
adjunction part of an EEC model can be equivalently given by a ‘state’ object of C (see [8,
Proof of Thm. 4] and Section 5.2).
Theorem 10.2. Every enriched model embeds in a closed enriched model.
Proof. The difference between enriched models and closed enriched models is that in an
enriched model (V, C) the value category V need not be cartesian closed nor have coproducts,
and the computation category C need not have coproducts, products and powers.
We use the Yoneda embedding to embed an enriched model in a closed enriched
model For any small category A we consider the category Ab of contravariant functors,
Aop → Set, and natural transformations between them. The Yoneda embedding A 7→
A(−, A) is a functor yA : A → Ab that exhibits Ab as a cocompletion of A. That is: Ab
is cocomplete, with colimits computed pointwise, and for any other cocomplete category
B and any functor F : A → B there is a colimit-preserving functor F! : Ab → B given by
def
π
F
F! (P ) = colim((yA ↓ P ) −
→A−
→ B), where (yA ↓ P ) is the category of elements of P ; this
colimit-preserving functor is essentially unique such that F ∼
= F! · yA .
b C)
b is a closed enriched model,
Let (V, C) be an enriched model. We will show that (V,
and that (V, C) embeds in it as an enriched model.

42

R. E. MØGELBERG AND S. STATON

We proceed by considering the following 2-categorical situation. Because the construction
d
(−) is a free cocompletion, it can be understood as a weak 2-functor from the 2-category
Cat of small categories, functors and natural transformations to the 2-category Cocomp of
categories with all colimits, colimit-preserving functors and natural transformations.
In fact it is necessary to be slightly more general than this: we will understand Cat and
Cocomp as 2-multicategories. Recall that a 2-multicategory is a Cat-enriched multicategory.
So it is like a 2-category except that the domains of the 1-cells are sequences of objects.
• The 2-multicategory Cat is defined as follows. The 0-cells are small categories with finite
coproducts. The 1-cells F : (A1 , . . . , An ) → B in Cat are functors in n arguments, i.e.
functors F : A1 × · · · × An → B. The 2-cells are natural transformations.
• The objects of the 2-multicategory Cocomp are categories with all colimits. The 1-cells
F : (A1 , . . . , An ) → B in Cocomp are functors F : A1 × · · · × An → B that preserve colimits
in each argument, i.e. that for fixed A1 ∈ A1 , . . . , An ∈ An and for 1 ≤ i ≤ n, the functor
F (A1 , . . . , −i , . . . An ) : Ai → B preserves colimits. The 2-cells are natural transformations.
d extends to a weak morphism of 2-multicategories from Cat to
• The construction (−)
Cocomp. A 1-cell F : (A1 , . . . , An ) → B in Cat is extended to a 1-cell in Cocomp,
c1 × · · · × A
cn → Bb which preserves colimits in each argument. This
i.e. a functor F! : A
construction is done by iteratively applying the following idea. If an n-ary functor
G : (A1 , . . . , Ak , . . . , An ) → B is such that B is cocomplete, A1 . . . Ak−1 are cocomplete, G
preserves colimits in each of the first (k − 1) arguments, and Ak is small, then there is an
ck , . . . , An ) → B that preserves colimits in each of the first k
n-ary functor G!k : (A1 , . . . , A
arguments such that G ∼
= G!k · (A1 , . . . , yA , . . . , An ). This is because the n-ary functor
G : (A1 , . . . , An ) → B can be curried to a functor
Ak → Cocomp(A1 , . . . , Ak−1 ; Cat(Ak+1 , . . . , An ; B))
whose codomain is cocomplete, and which can thus be extended to a colimit-preserving
functor using the universal property of Âk :
Âk → Cocomp(A1 , . . . , Ak−1 ; Cat(Ak+1 , . . . , An ; B));
ck , . . . , An ) → B. Ultimately, the extension
this can be uncurried to give G!k : (A1 , . . . , A
c1 × · · · × A
cn → Bb satisfies the following coend formula:
F! : A
Z A1 ,...,An
∼
F! (P1 , . . . , Pn )(B) =
P1 (A1 ) × · · · × Pn (An ) × B(B, F (A1 , . . . , An ))
• Recall the following consequence of the special adjoint functor theorem: a morphism
c1 , . . . , A
cn ) → B in Cocomp can be equivalently described as a functor that has a
F : (A
right adjoint in each argument, i.e. a right adjoint for each functor F (P1 , . . . , −i , . . . , Pn ) :
ci → B.
A
• Aside from size issues, there is a forgetful morphism of 2-multicategories Cat → Cocomp
d is left biadjoint to it. The Yoneda embedding is the unit for this adjunction.
and (−)
With the general situation explained, the proof of Theorem 10.2 is straightforward. We
begin by considering the evident notion of ‘weak monoid’ in a 2-multicategory K. This
comprises an object M of K and two 1-cells: m : (M, M) → M and e : () → M, with three
coherence 2-isomorphisms. A morphism of 2-multicategories K → K0 takes weak monoids
in K to weak monoids in K0 . In particular a monoidal category is a weak monoid in Cat, and
d takes it to a weak monoid in Cocomp, which is a cocomplete biclosed
the construction (−)

LINEAR USAGE OF STATE

43

c preserves the weak monoid structure.
monoidal category. The Yoneda embedding M → M
This is Day’s convolution construction [5, 17].
In particular, the value category V of our enriched model has products and this exhibits
b is cartesian closed and that the Yoneda
it as a monoidal category. It follows that V
b preserves the product structure in V. Given a weak monoid M in a
embedding V → V
2-multicategory K, we consider the evident notion of weak action for M: an object A of K
and a 1-cell (M, A) → A satisfying the laws of monoid actions up-to coherent isomorphism.
A morphism of 2-multicategories takes weak monoid actions to weak monoid actions. In
particular, given a monoidal category M, an action of M on another category A induces an
c with powers and copowers. The Yoneda embedding A → Ab preserves
enrichment of Ab in M
the monoidal action. Moreover since it is 2-natural it preserves any enrichment or powers
that already exist in A.
b is enriched in V
b with
In particular, in our enriched model, V acts on C and so C
b
powers and copowers, and the Yoneda embedding C → C is enriched in V and preserves
copowers.
The crux of the proof is that the Yoneda embedding adds closed structure — cartesian
closed structure and powers — while preserving the other structure. Although the Yoneda
embedding does not freely add the closed structure, it is considerably simpler that the free
closure. This is one reason why Yoneda embeddings are a common technique in semantics.
For instance, the enriched Yoneda embedding is used by Egger et al. [8] to show that
Levy’s call-by-push-value embeds in the enriched effect calculus (but the enriched Yoneda
embedding is not appropriate in our proof because it does not preserve copowers).
Theorem 10.2 explains that EEC is conservative over ECBV, but it neglects sum types.
Sum types play an important role in the study of generic effects and state access operations
(§8). We now show that EEC is conservative over ECBV with sum types.
Proposition 10.3. Every distributive enriched model (§6.1) embeds in a closed enriched
model (Def. 10.1).
Proof. Our proof of Proposition 10.3 follows the same outline as our proof of Theorem 10.2.
We must modify that proof because the Yoneda embedding yA : A → Ab does not preserve
coproducts. We use the following variation on the Yoneda embedding. For any category A
with finite coproducts, let FP(Aop , Set) be the category of finite-product-preserving functors
Aop → Set and natural transformations between them. Assuming A has coproducts, the
category FP(Aop , Set) has coproducts too. The Yoneda embedding A 7→ A(−, A) is a functor
A → FP(Aop , Set) which preserves coproducts. In fact, the Yoneda embedding exhibits
FP(Aop , Set) as the cocompletion of A as a category with finite coproducts. The category
FP(Aop , Set) is cocomplete (although not all colimits are computed pointwise), and for any
coproduct-preserving functor F : A → B into a cocomplete category B there is a colimitdef
π
F
preserving functor F! : FP(Aop , Set) → B given by F! (P ) = colim((yA ↓ P ) −
→A−
→ B);
this colimit-preserving functor is essentially unique such that F ∼
= F! · yA (see e.g. [21,
Thms 5.86, 6.11], [42], [11]). Since a distributive enriched model has coproducts, this is the
right variation of the Yoneda embedding to use.
We now mimic the proof of Theorem 10.2, replacing the cocompletion construction
d
(−) with the cocompletion FP((−)op , Set) of a category with coproducts. Consider the
2-multicategory Coprod: the 0-cells are small categories with finite coproducts; the 1-cells
F : (A1 , . . . , An ) → B are functors F : A1 × · · · × An → B that preserve coproducts in each

44

R. E. MØGELBERG AND S. STATON

argument; the 2-cells are natural transformations. The construction FP((−)op , Set) extends
to a morphism of 2-multicategories from Coprod to Cocomp. By the special adjoint functor
op
theorem, a morphism (FP(Aop
1 , Set), . . . , FP(An , Set)) → B in Cocomp is a functor that
has a right adjoint in each argument.
A weak monoid M in Coprod is a distributive monoidal category, i.e., a monoidal
category with coproducts such that the tensor preserves coproducts in each argument. The
construction FP(−op , Set) takes it to a weak monoid in Cocomp, which is a cocomplete
biclosed monoidal category. The Yoneda embedding M → FP(Mop , Set) preserves the
weak monoid structure and coproducts.
In particular, if (V, C) is a distributive enriched model then V has distributive products
and this exhibits it as a distributive monoidal category. It follows that FP(Vop , Set)
is cartesian closed with coproducts and that the Yoneda embedding V → FP(Vop , Set)
preserves the coproduct and product structure in V.
An action of a weak monoid in Coprod is the same thing as a distributive action in the
sense of Section 6.1. Given a distributive monoidal category M, a distributive action of M on
a category A with finite coproducts induces an enrichment of FP(Aop , Set) in FP(Mop , Set)
with powers and copowers. The Yoneda embedding A → FP(Aop , Set) preserves coproducts
and the monoidal action. Moreover since it is 2-natural it preserves any enrichment or
powers that already exist in A.
In particular, if (V, C) is a distributive enriched model then V acts on C, and so
FP(Cop , Set) is enriched in FP(Vop , Set) with powers and copowers, and the Yoneda
embedding C → FP(Cop , Set) is enriched in V and preserves coproducts and copowers.
The construction in this proof is related to the following natural situation. Let Setf be
the category of finite sets, and let T be a Lawvere theory. Then (Setf , Top ) is almost an
enriched model, except that the category Top is typically not Setf -enriched. Nonetheless, our
construction applied to (Setf , Top ) yields the basic motivating example of an EEC model:
FP(Setop
f , Set) is the category of sets (since Setf is the free category with finite coproducts
on one generator) and FP(T, Set) is the category of algebras of the Lawvere theory. (See
also [41, Thm. 38].)
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Appendix A. Categories of models
A.1. The 2-category Enr of enriched models. We first define a notion of morphism of
enriched call-by-value model and transformations between these. This gives a 2-category
Enr. Let (V, C) and (V0 , C0 ) be enriched call-by-value models (Def. 2.1). A morphism from
(V, C) to (V0 , C0 ) is a triple (F, G, λ) such that F : V → V0 and G : C → C0 are functors,
and λ is a natural family of isomorphisms
λA,B : G(A · B) ∼
= F (A) · G(B)
The following three conditions must be satisfied:
• F preserves products (up to isomorphism).
• The following two coherence diagrams commute
λ
G(1 · B) - F 1 · GB
∼
=

! ·S

∼
=
GB 
?

G((A × B) · C)

∼
=

(A.1)

?

1 · GB
λ
-

- G(A · B · C)

F A · G(B · C)
FA · λ

λ
?

F (A × B) · GC

∼
=
-

hF π1 , F π2 i · GC
-

?

(F A × F B) · GC
F A · F B · GC
0
∼ C (GB, GC)
• The mate of
is an isomorphism F (C(B, C)) =
−1
Recall that the mate of λ is the adjoint correspondent of
λ−1

F (C(B, C)) · GB

λ−1
-

G(C(B, C) · B)

G(ev)

where ev is the counit of the adjunction (−) · B a C(B, (−)).

G(C)

(A.2)

LINEAR USAGE OF STATE

47

The 2-cells between morphisms (F, G, λ), (F 0 , G0 , λ0 ) : (V, C) → (V0 , C0 ) are natural
isomorphisms β : F ∼
= F 0, γ : G ∼
= G0 making the following diagram commute:
λA,B
- F A · GB
G(A · B)
β·γ

γ
?

G0 (A · B)

λ0A,B

(A.3)

?

- F 0 A · G0 B

The composition of 1-cells is defined as (F 0 , G0 , λ0 ) ◦ (F, G, λ) = (F 0 F, G0 G, λ0 ◦ G0 λ):
G0 G(A · B)

G0λ

λ0
G0 (F A · GB) - F 0 F A · G0 GB

The composition of 2-cells is simply pointwise (β 0 , γ 0 ) ◦ (β, γ) = (β 0 ◦ β, γ 0 ◦ γ)
A.2. The 2-category Ecbv. The objects of the 2-category Ecbv are tuples (V, C, S) where
(V, C) is an enriched call-by-value model and S is a ‘state’ object in C. A 1-cell (V, C, S) →
(V0 , C0 , S 0 ) is a quadruple (F, G, λ, δ) such that (F, G, λ) : (V, C) → (V0 , C0 ) is a morphism
of enriched call-by-value models and δ is an isomorphism
δ : GS ∼
= S0
A 2-cell from (F, G, λ, δ) to (F 0 , G0 , λ0 , δ 0 ) is a 2-cell of enriched call-by-value models (§A.1)
(β, γ) : (F, G) → (F 0 , G0 )
such that
GS

δ
0

δ

γ

- S0
-

(A.4)

?

G0 S
Composition of 1-cells and 2-cells in Ecbv is defined as in Enr. The state object
isomorphisms are composed as follows:
G0 δ - 0 0
δ 0 - 00
GS
S .
G0 GS
A.3. The 2-category Kleisli. We introduce the 2-category Kleisli whose objects are enriched
Kleisli models (§3.1). Morphisms from (V, C, J) to (V0 , C0 , J 0 ) are morphisms of enriched
models (F, G, λ) : (V, C) → (V0 , C0 ) such that J 0 ◦ F = G ◦ J. Note in particular that this
means that F and G agree on objects, since J and J 0 are required to be identities on objects.
A 2-cell (F, G, λ) → (F 0 , G0 , λ0 ) is a 2-cell of enriched models (β, γ) : (F, G) → (F 0 , G0 ) such
that γJ = J 0 β : GJ → J 0 F 0 .
A.4. Sums. The category dEnr is defined to have distributive enriched models as objects,
1-cells as in Enr with the restriction that F and G must both preserve coproducts, and
2-cells as in Enr. The definitions of the 2-categories Ecbv and Kleisli extend to 2-categories
dEcbv and dKleisli of distributive enriched models with state objects and distributive Kleisli
models.

48

R. E. MØGELBERG AND S. STATON

A.5. Effect theories. We now explain how to extend the 2-categories to models that
support effect theories, expanding on Section 8.8. We first define what it means for a functor
to preserve a value theory.
Definition A.1. Let V and V0 be distributive categories with given interpretations of some
fixed value theory (§8.5.1). A functor F : V → V0 preserving products and coproducts
preserves the interpretation of the value theory if F [[α]] = [[α]] for all type constants α of the
theory, and the diagram
F [[f-]]
F ([[ᾱ]])
F ([[β]])
]
]
[[f
hF π1 , . . . , F πn i
?

F [[α1 ]] × · · · × F [[αn ]]
commutes for all term constants f : ᾱ → β of the theory.
A.5.1. The 2-category dEcbvE . For any effect theory E, the objects of the 2-category dEcbvE
are distributive enriched models with state and a chosen E-comodel structure on the state
object (§8.5.3). Morphisms (1-cells) are morphisms (F, G, λ, δ) of dEcbv such that
• F preserves the value theory of E as in Definition A.1.
• The comodel structure is preserved, i.e., for each effect constant e : β̄; ᾱ1 + . . . + ᾱn of E,
the following diagram commutes
G[[e]]
- G(([[ᾱ1 ]] + · · · + [[ᾱn ]]) · S)
G([[β̄]] · S)
(A.5)
[[e]] [[β̄]] · S 0
([[ᾱ1 ]] + · · · + [[ᾱn ]]) · S 0
where the vertical maps are constructed using λ, δ and the preservation of products and
coproducts. (This makes sense because F ([[α]]) = [[α]] for all type constants α.)
A 2-cell of dEcbvE is a 2-cell (β, γ) : (F, G, δ) → (F 0 , G0 , δ 0 ) of dEcbv such that β[[α]] is the
identity for all type constants α of the theory E.
?

?

A.5.2. The 2-category dKleisliE . For any effect theory E, the objects of the 2-category
dKleisliE are distributive Kleisli models with given interpretations of E. Morphisms are
morphisms of dKleisli such that
• F preserves the value theory of E as in Definition A.1
• The interpretation of effect constants is preserved, i.e., for each e : β̄; ᾱ1 + . . . + ᾱn of E,
the following diagram commutes
G[[e]]
- G([[ᾱ1 ]] + · · · + [[ᾱn ]])
G([[β̄]])
hGJ(πi )ii

(A.6)
?

[[β̄]]

[[e]]-

?

([[ᾱ1 ]] + · · · + [[ᾱn ]])

where the left vertical map types because G[[βi ]] = [[βi ]], and the right vertical map is
constructed using the fact that G preserves coproducts.

LINEAR USAGE OF STATE

49

The 2-cells of dKleisliE are the 2-cells (β, γ) of dKleisli with β[[α]] the identity for all type
constants α of the theory E.
Appendix B. Bi-initiality of syntactic models
We sketch a proof of Theorem 2.4: the syntactic enriched model is bi-initial in the 2-category
Ecbv of enriched models. Bi-initiality of the syntactic monad model (Theorem 3.3) can be
proved in a similar manner.
Lemma B.1. Suppose (F, G, δ) : (V, C, S) → (V0 , C0 , S 0 ) is a morphism in dEcbvE . Let
[[−]] be the interpretation of ECBV in (V, C), and let [[−]]0 be the interpretation of ECBV in
(V0 , C0 ) (as in §2.2, §8.5.3). The family of isomorphisms given by
id : [[α]]0 → F [[α]]
δ −1 : [[S]]0 = S 0 → GS = G[[S]]
extends uniquely to a type-indexed family of isomorphisms [[A]]0 ∼
= F ([[A]]) and [[B]]0 ∼
= G([[B]])
such that if Γ | − ` t : A and Γ | ∆ ` u : B then
[[Γ]]0

[[t]]0 -

∼
=

[[A]]0

and

∼
=
?

F [[Γ]]

F [[t]]-

[[u]]0 -

[[Γ]]0 · [[∆]]0
∼
=

?

Moreover, if (β, γ) is a 2-cell then
∼
=[[A]]0
F [[A]]
=∼
β
-

∼
=
?

G([[Γ]] · [[∆]])

F [[A]]

?

F 0 [[A]]

and

[[B]]0

[[B]]0

G[[u]]-

∼
=-

(B.1)
?

G[[B]]

G[[B]]

=∼

γ
-

(B.2)

?

G0 [[B]]

Proof notes. The isomorphisms are defined by induction on A and B. For example, in the
case of !A ⊗ B we use the composite
∼
λ−1
=
[[!A ⊗ B]]0 = [[A]]0 · [[B]]0 - (F [[A]]) · (G[[B]]) - G([[A]] · [[B]]) = G[[!A ⊗ B]]
Commutativity of (B.1) and (B.2) is proved by induction on structure of terms and types
respectively. We omit the lengthy verifications entirely, but remark that (A.1) is used to
prove the case of linear variable introduction, and diagram (A.2) is used to prove the case of
linear function application and the elimination rule for !A ⊗ B. The case of effect constants
is exactly the requirement (A.5).
Uniqueness is proved by induction on types.
Bi-initiality of the syntactic model (Theorem 2.4) follows from Lemma B.1 as follows.
Given any other model, the unique morphism is given by interpretation of the syntax. To
prove uniqueness, suppose (F, G, δ) is a 1-cell from dEcbvE to (V, C, S). Then Lemma B.1
gives the natural isomorphism: since interpretation in the syntactic model of terms with one
variable is simply the identity, diagrams (A.1) prove naturality of the isomorphism. The
required commutative diagrams (A.3) for 2-cells follow directly from definitions.

50

R. E. MØGELBERG AND S. STATON

Appendix C. The adjunction between Kleisli models and enriched models
C.1. The 2-functor St : Kleisli → Ecbv. We define the 2-functor St : Kleisli → Ecbv by
the following data:
def

St(V, C, J) = (V, C, 1)
def

St(F, G) = (F, G, J 0 (! ))

for (F, G) : (V, C, J) → (V0 , C0 , J 0 )

def

St(β, γ) = (β, γ)
Note that J 0 (! ) has the right type:
J 0 (! )

G(1) = G(J(1)) = J 0 (F (1)) −−−→ J 0 (1) = 1.
C.2. The 2-functor Kl : Ecbv → Kleisli. Recall that Kl is defined on objects in Section 5
as
def

Kl(V, C, S) = (V, KlS , JS )
where the category KlS has the same objects as V and homsets
def

HomKlS (A, B) = HomC (A · S, B · S)
def

On morphisms we define Kl(F, G, λ) = (F, Kl(F, G), λKl ) where
Kl(F, G) : KlS → KlS 0
is the functor that maps an object A to F A and a morphism f : A · S → B · S to the following
composite:
F A · S0

A · δ-

F A · GS

λ−1
-

G(A · S)
G(f )

Kl(F, G)(f )
B·
F B · S0 
?

δ −1

(C.1)

λ
F B · GS  G(B · S)
?

The natural transformation
λKl : Kl(F, G)(A ·Kl B) → F (A) ·Kl Kl(F, G)(B)
has components given by
hF π1 , F π2 i · S 0 : F (A × B) · S 0 → (F (A) × F (B)) · S 0
On 2-cells we define Kl(β, γ) = (β, Kl(β, γ)) where Kl(β, γ) is the natural transformation
whose components are Kl(β, γ)A = βA · S 0 : F A · S 0 → F 0 A · S 0 . Note that this defines a
morphism in KlS 0 from Kl(F, G)(A) = F A to Kl((F 0 , G0 ))(A) = F 0 A as required.

LINEAR USAGE OF STATE

51

C.3. The unit. If (V, C, J) is a distributive Kleisli model, then
Kl(St(V, C, J)) = (V, Kl1 , J1 )
where Kl1 has the same objects as V and HomKl1 (A, B) = HomC (A × 1, B × 1) and
J1 (A) = A, J1 (f ) = J(f × 1) = f · 1.
We define the unit η(V,C,J) as
def

η(V,C,J) = (id , H, id ) : (V, C, J) → (V, Kl1 , J1 )
where H(A) = A, and H(f : A → B) is the composition
A×1

J(π1 ) -

f

A

- B

Jhid , !-i

B×1

The third component of η should be a natural transformation H(A · B) → A ·Kl HB. Since
both sides are equal to A × B it makes sense to take the identity transformation.
Lemma C.1. Each η(V,C,J) is an isomorphism in Kleisli.
The next lemma states that the unit is a 2-natural transformation.
Lemma C.2. For each pair of distributive Kleisli models (V, C, J) and (V0 , C0 , J 0 ), the
following diagram of functors commutes.
Kleisli((V, C, J), (V0 , C0 , J 0 ))

Kl◦St

/ Kleisli((V, Kl1 , J1 ), (V0 , Kl1 0 , J 0 ))
1
-



(η(V0 ,C0 ,J 0 ) )∗

(η(V,C,J) )∗

Kleisli((V, C, J), (V0 , Kl1 0 , J1 0 ))

where f ∗ and g∗ are the functors given by pre- and postcomposition by f and g respectively.
C.4. The counit. Let (V, C, S) be a model of enriched call-by-value. Then
St(Kl(V, C, S)) = (V, KlS , 1)
Recall that KlS has the same objects as V and set of morphisms defined as
HomKlS (A, B) = HomC (A · S, B · S)
Define the counit  as
def

(V,C,S) = (id , IS , δ) : (V, KlS , 1) → (V, C, S)
where IS is the functor defined by IS (A) = A · S, IS (f ) = f , and δ is the isomorphism
δ : IS (1) = 1 · S → S
The counit is not strictly natural, but it does satisfy the following naturality condition.
Lemma C.3. Let (V, C, S), (V0 , C0 , S 0 ) be two given objects of Ecbv. The following diagram
commutes up to natural isomorphism.
Ecbv((V, C, S), (V0 , C0 , S 0 ))

St◦Kl

/ Ecbv((V, KlS , 1), (V0 , Kl 0 , 1))
S

∼
=
((V,C,S) )∗

-



((V0 ,C0 ,S 0 ) )∗

Ecbv((V, KlS , 1), (V0 , C0 , S 0 ))

52

R. E. MØGELBERG AND S. STATON

C.5. Triangle equalities. One of the triangle equalities only holds up to isomorphism.
Proposition C.4. Let (V, C, S) be a distributive enriched model with state. The composite
1-cell
Kl((V,C,S) ) ◦ ηKl(V,C,S)
is the identity on Kl(V, C, S).
Proposition C.5. Let (V, C, J) be a distributive Kleisli model. The composite 1-cell
St(V,C,J) ◦ St(η(V,C,J) )
is naturally isomorphic to the identity on St(V, C, J).
C.6. Proof of Theorem 5.3. We now prove Theorem 5.3: η and  induce an equivalence
of categories
Ecbv(St(V, C, J), (V0 , C0 , S)) ' Kleisli((V, C, J), Kl(V0 , C0 , S)).
In the following we use X to denote an object of Ecbv (rather than the much longer
(V, C, S)) and Y to denote an object of Kleisli. The required equivalence of categories
consists of the composite functors
∗
Kl
(ηX )Kleisli(X, Kl(Y ))
Ecbv(St(X), Y ) - Kleisli(Kl(St(X)), Kl(Y ))
St
(Y )∗Ecbv(St(X), Y )
Kleisli(X, Kl(Y )) - Ecbv(St(X), St(Kl(Y )))
We need to prove that the two composites of these are naturally isomorphic to identities.
For the first of the composites, consider the following sequences of identities and natural
isomorphisms
∗
(Y )∗ ◦ St ◦ ηX
◦ Kl = (Y )∗ ◦ (St(ηX ))∗ ◦ St ◦ Kl

= (St(ηX ))∗ ◦ (Y )∗ ◦ St ◦ Kl
∼
= (St(ηX ))∗ ◦ (St(Y ) )∗

Lemma C.3

∗

= (St(Y ) ◦ St(ηX ))
∼ id ∗
=

Proposition C.5

= id
The other composite is isomorphic to the identity for a similar reason.

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