Functional Programming

If there is one structure that permeates category theory and, by implication, the whole of mathematics, it’s the monoid. To study the evolution of this concept is to study the power of abstraction and the idea of getting more for less, which is at the core of mathematics. When I say “evolution” I don’t necessarily mean chronological development. I’m looking at a monoid as if it were a life form evolving through various eons of abstraction.

It’s an ambitious project and I’ll have to cover a lot of material. I’ll start slowly, with the definitions of magmas and monoids, but then I will accelerate. A lot of concepts will be introduced in one or two sentences, mainly to familiarize the reader with the notation. I’ll dwell a little on monoidal categories, then breeze through ends, coends, and profunctors. I’ll show you how monads, arrows, and applicative functors arise from monoids in various monoidal categories.


The Magmas of the Hadean Eon

Monoids evolved from more primitive life forms feeding on sets. So, before even touching upon monoids, let’s talk about cartesian products, relations, and functions. You take two sets a and b (or, in the simplest case, two copies of the same set a) and form pairs of elements. That gives you a set of pairs, a.k.a., the cartesian product a×b. Any subset of such a cartesian product is called a relation. Two elements x and y are in a relation if the pair <x, y> is a member of that subset.

A function from a to b is a special kind of relation, in which every element x in the set a has one and only one element y in the set b that’s related to it. (Sometimes this is called a total function, since it’s defined for all elements of a).

Even before there were monoids, there was magma. A magma is a set with a binary operation and nothing else. So, in particular, there is no assumption of associativity, and there is no unit. A binary operation is simply a function from the cartesian product of a with itself back to a

a × a -> a

It takes a pair of elements <x, y>, both coming from the set a, and maps it to an element of a.

It’s tempting to quote the Haskell definition of a magma:

class Magma a where
  (<>) :: a -> a -> a

but this definition is already tainted with some higher concepts like currying. An alternative would be:

class Magma a where
  (<>) :: (a, a) -> a

Here, we at least see a pair of elements that are being “multiplied.” But the pair type (a, a) is also a higher-level concept. I’ll come back to it later.

Lack of associativity means that we cannot identify (x<>y)<>z with x<>(y<>z). You have to keep the parentheses.

You might have heard of quaternions — their multiplication is associative. But not many people have heard of octonions, which are not associative. In fact Hamilton, who discovered quaternions, invented the word associative to disassociate himself from octonions, which are not.

If you’re familiar with continuous groups, you might know that Lie algebras are not associative.

Closer to home — most operations on floating-point numbers are not associative on modern computers because of rounding errors.

But, really, most interesting binary operations are associative. So out of the magma emerges a semigroup. In a semigroup you can drop parentheses. A non-trivial (that is, non-monoidal) example of a semigroup is the set of integers with max binary operation. A maximum of three numbers is the same no matter in which order you pair them. But there is no integer that’s less or equal to any other integer, so this is not a monoid.

Monoids of the Archean Eon

But, really, most interesting binary operations are both associative and unital. There usually is a “do nothing” element with respect to most binary operations. So life as we know it begins with a monoid.

A monoid is a set with a binary operation that is associative, and with a special element called the unit e that is neutral with respect to the binary operation. To be precise, these are the three monoid laws:

(x <> y) <> z = x <> (y <> z)
e <> x = x
x <> e = x

In Haskell, the traditional definition of a monoid uses mempty for the unit and mappend for the binary operation:

class Monoid a where
    mempty  :: a
    mappend :: a -> a -> a

As with the magma, the definition of mappend is curried. Equivalently, it could have been written as:

mappend :: (a, a) -> a

I’ll come back to this point later.

There are plenty of examples of monoids. Non-negative integers with addition, or positive integers with multiplication are the obvious ones. Strings with concatenation are interesting too, because concatenation is not commutative.

Just like pairs of elements from two sets a and b organize themselves into a set a×b, which is their cartesian product; functions between two sets organize themselves into a set — the set of functions from a to b, which we sometimes write as a->b.

This organizing principle is characteristic of sets, where everything you can think of is a set. Except when it’s more than just a set — for instance when you try to organize all sets into one large collection. This collection, or “class,” is not itself a set. You can’t have a set of all sets, but you can have a category Set of “small” sets, which are sets that belong to a “universe.” In what follows, I will confine myself to a single universe in order to dodge questions from foundational mathematicians.

Let’s now pop one level up and look at cartesian product as an operation on sets. For any two sets a and b, we can construct the set a×b. If we view this as “multiplication” of sets, we can say that sets form a magma. But do they form a monoid? Not exactly! To begin with, cartesian product is not associative. We can see it in Haskell: the type ((a, b), c) is not the same as the type (a, (b, c)). They are, however, isomorphic. There is an invertible function called the associator, from one type to the other:

alpha :: ((a, b), c) -> (a, (b, c))
alpha ((x, y), z) = (x, (y, z))

It’s just a repackaging of containers (such repackaging is, by the way, called a natural transformation).

For the unit of this “multiplication” we can pick the singleton set. In Haskell, this is the type called unit and it’s denoted by an empty pair of parentheses (). Again, the unit laws are valid up to isomorphism. There are two such isomorphisms called left and right unitors:

lambda :: ((), a) -> a
lambda ((), x) = x
rho :: (a, ()) -> a
rho (x, ()) -> x

We have just exposed monoidal structure in the category Set. Set is not strictly a monoid because monoidal laws are satisfied only up to isomorphism.

There is another monoidal structure in Set. Just like cartesian product resembles multiplication, there is an operation on sets that resembles addition. It’s called disjoint sum. In Haskell it’s embodied in the type Either a b . Just like cartesian product, disjoint sum is associative up to isomorphism. The unit (or the “zero”) of this sum type is the empty set or, in Haskell, the Void type — also up to isomorphism.

The Cambrian Explosion of Categories

The first rule of abstraction is, You do not talk about Fight Club. In the category Set, for instance, we are not supposed to admit that sets have elements. An object in Set is really a set, but you never talk about its elements. We still have functions between sets, but they become abstract morphisms, of which we only know how they compose.

Composition of functions is associative, and there is an identity function for every set, which serves as a unit of composition. We can write these rules compactly as:

(f ∘ g) ∘ h = f ∘ (g ∘ h)
id ∘ f = f
f ∘ id = f

These look exactly like monoid laws. So do functions form a monoid with respect to composition? Not quite, because you can’t compose any two functions. They must be composable, which means their endpoints have to match. In Haskell, we can compose g after f, or g ∘ f, only if:

f :: a -> b
g :: b -> c

Also, there is no single identity function, but a whole family of functions ida, one for each set a. In Haskell, we call that a polymorphic function.

But notice what happens if we restrict ourselves to just a single object a in Set. Every morphism from a back to a can be composed with any other such morphism (their endpoints always match). Moreover, we are guaranteed that among those so called endomorphisms there is one identity morphism ida, which acts as a unit of composition.

Notice that I switched from the set/function nomenclature to the more general object/morphism naming convention of category theory. We can now forget about sets and functions and define an arbitrary category as a collection (a set in a given universe) of objects, and sets of morphisms that go between them. The only requirements are that any two composable morphisms compose, and that there is an identity morphism for every object. And that composition must be associative.

We can now forget about sets and define a monoid as a category that has only one object. The binary operation is just the composition of (endo-)morphisms. It works! We have defined a monoid without a set. Or have we?

No, we haven’t! We have just swept it under the rug — the rug being the set of morphisms. Yes, morphisms between any two objects form a set called the hom-set. In a category C, the hom-set between objects a and b is denoted by C(a, b). So we haven’t completely eliminated sets from the picture.

In the single object category M, we have only one hom-set M(a, a). The elements of this set — and we are allowed to call them elements because it’s a set — are morphisms like f and g. We can compose them, and we can call this composition “multiplication,” thus recovering our previous definition of the monoid as a set. We get associativity for free, and we have the identity morphism ida serving as the unit.

It might seem at first that we haven’t made progress and, in fact, we might have made some things more complicated by forgetting the internal structure of objects. For instance, in the category Set, it’s no longer obvious what an empty set is. You can’t say it’s a set with no elements because of the Fight Club rule. Similarly with the singleton set. Fortunately, it turns out that both these sets can be uniquely described in terms of their interactions with other sets. By that I mean the kind of functions/morphisms that connect them to other objects in Set. These object-opaque definitions are called universal constructions. For instance, the empty set is the only set that has a unique morphism going from it to every other set. The advantage of this characterization is that it can now be applied to any category. One may ask this question in any category: Is there an object that has this property? If there is, we call it the initial object. The empty set is the initial object in Set. Similarly, a singleton set is the terminal object in Set (and it’s unique up to unique isomorphism).

A cartesian product of two sets can also be defined using a universal construction, one which doesn’t mention elements (or pairs of elements). And again, this construction may be used to define a (categorical) product in other categories. Of particular interest are categories where a product exists for every pair of objects (it does in Set).

In such categories there is actually an even better way of defining a product using an adjunction. But before we can get to adjunctions, let me summarize a few millions of years of evolution in a few terse paragraphs.

A functor is a mapping of categories that preserves their structure. It maps objects to objects and morphisms to morphisms. In Haskell we define a functor (really, an endofunctor) as a type constructor f (a mapping from types to types) that can be lifted to functions that go between these types:

class Functor f where
  fmap :: (a -> b) -> (f a -> f b)

The mapping of morphisms must also preserve composition and identity. Functors may collapse multiple objects into one, and multiple morphisms into one, but they never break connections. You may also think of functors as embedding one category inside another.

Finally, functors can be composed in the obvious way, and there is an identity endofunctor that maps a category onto itself. It follows that categories (at least the small ones) form a category Cat in which functors serve as morphisms.

There may be many ways of embedding one category inside another, and it’s extremely useful to be able to compare such embeddings by defining mappings between them. If we have two functors F and G between two categories C and D we define a natural transformation between these functors by picking a morphism between a pair F a and G a, for every a.

In Haskell, a natural transformation between two functors f and g is a polymorphic function:

type Nat f g = forall a. f a -> g a

In general, natural transformations must obey additional naturality conditions, but in Haskell they come for free (due to parametricity).

Natural transformations may be composed, and there is an identity natural transformations from any functor to itself. It follows that functors between any two categories C and D form a category denoted by [C, D], where natural transformations play the role of morphisms. A hom-set in such a category is a set of natural transformations between two functors F and G denoted by [C, D](F, G).

An invertible natural transformation is called a natural isomorphism. If two functors are naturally isomorphic they are essentially the same.

Arthropods and their Adjoints

Using a pair of functors that are the inverse of each other we may define equivalence of categories, but there is an even more useful concept of adjoint functors that compare the structures of two non-equivalent categories. The idea is that we have a “right” functor R going from category C to D and a “left” functor L going in the other direction, from D to C.

Adj - 1

There are two possible compositions of these functors, both resulting in round trips or endofunctors. The categories would be equivalent if those endofunctors were naturally isomorphic to identity endofunctors. But for an adjunction, we impose weaker conditions. We require that there be two natural transformations (not necessarily isomorphisms):

η :: ID -> R ∘ L
ε :: L ∘ R -> IC

The first transformation η is called the unit; and the second ε, the counit of the adjunction.

In a small category objects form sets, so it’s possible to form a cartesian product of two small categories C and D. Object in such a category C×D are pairs of objects <c, d>, and morphisms are pairs of morphisms <f, g>.

After these preliminaries, we are ready to define the categorical product in C using an adjunction. We chose C×C as the left category. The left functor is the diagonal functor Δ that maps any object c to a pair <c, c> and any morphism f to a pair of morphisms <f, f>. Its right adjoint, if it exists, maps a pair of objects <a, b> to their categorical product a×b.


Interestingly, the terminal object can also be defined using an adjunction. This time we chose, as the left category, a singleton category with one object and one (identity) morphism. The left functor maps any object c to the singleton object. Its right adjoint, if it exists, maps the singleton object to the terminal object in C.

A category with all products and the terminal object is called a cartesian category, or cartesian monoidal category. Why monoidal? Because the operation of taking the categorical product is monoidal. It’s associative, up to isomorphism; and its unit is the terminal object.

Incidentally, this is the same monoidal structure that we’ve seen in Set, but now it’s generalized to the level of other categories. There was another monoidal structure in Set induced by the disjoint sum. Its categorical generalization is given by the coproduct, with the initial object playing the role of the unit.

But what about the set of morphisms? In Set, morphisms between two sets a and b form a hom-set, which is the object of the same category Set. In an arbitrary category C, a hom-set C(a, b) is still a set — but now it’s not an object of C. That’s why it’s called the external hom-set. However, there are categories in which each external hom-set has a corresponding object called the internal hom. This object is also called an exponential, ba. It can be defined using an adjunction, but only if the category supports products. It’s an adjunction in which the left and right categories are the same. The left endofunctor takes an object b and maps it to a product b×a, where a is an arbitrary fixed object. Its adjoint functor maps an object b to the exponential ba. The counit of this adjunction:

ε :: ba × a -> b

is the evaluation function. In Haskell it has the following signature:

eval :: (a -> b, a) -> b

The Haskell function type a->b is equivalent to the exponential ba.

A category that has all products and exponentials together with the terminal object is called cartesian closed. Cartesian closed categories, or CCCs, play an important role in the semantics of programming languages.

Tensorosaurus Rex

We have already seen two very similar monoidal structures induced by products and coproducts. In mathematics, two is a crowd, so let’s look for a pattern. Both product and coproduct act as bifunctors C×C->C. Let’s call such a bifunctor a tensor product and write it as an infix operator a ⊗ b. As a bifunctor, the tensor product can also lift pairs of morphisms:

f :: a -> a'
g :: b -> b'
f ⊗ g :: a ⊗ b -> a' ⊗ b'

To define a monoid on top of a tensor product, we will require that it be associative — up to isomorphism:

α :: (a ⊗ b) ⊗ c -> a ⊗ (b ⊗ c)

We also need a unit object, which we will call i. The two unit laws are:

λ :: i ⊗ a -> a
ρ :: a ⊗ i -> a

A category with a tensor product that satisfies the above properties, plus some additional coherence conditions, is called a monoidal category.

We can now specialize the tensor product to categorical product, in which case the unit object is the terminal object; or to coproduct, in which case we chose the initial object as the unit. But there is an even more interesting operation that has all the properties of the tensor product. I’m talking about functor composition.



Functors between any two categories C and D form a functor category [C, D] with natural transformations playing the role of morphisms. In general, these functors don’t compose (their endpoints don’t match) unless we pick the target category to be the same as the source category.

Endofunctor Composition

In the endofunctor category [C, C] any two functors can be composed. But in [C, C] functors are objects, so functor composition becomes an operation on objects. For any two endofunctors F and G it produces a new endofunctor F∘G. It’s a binary operation, so it’s a potential candidate for a tensor product. Indeed, it is a bifunctor: it can be lifted to natural transformations, which are morphisms in [C, C]. It’s associative — in fact it’s strictly associative, the associator α is the identity natural transformation. The unit with respect to endofunctor composition is the identity functor I. So the category of endofunctors is a monoidal category.

Unlike product and coproduct, which are symmetric up to isomorphism, endofunctor composition is not symmetric. In general, there is no relation between F∘G and G∘F.

Profunctor Composition

Different species of functors came up with their own composition strategies. Take for instance the profunctors, which are functors Cop×D->Set. They generalize the idea of relations between objects in C and D. The sets they map to may be thought of as sets of proofs of the relationship. An empty set means that the two objects are not related. If you want to compose two relations, you have to find an element that’s common to the image of one relation and the source of the other (relations are not, in general, symmetric). The proofs of the new composite relation are pairs of proofs of individual relations. Symbolically, if p and q are such profunctors/relations, their composition can be written as:

exists x. (p a x, q x b)

Existential quantification in Haskell translates to polymorphic construction, so the actual definition is:

data PCompose p q a b = forall x . PCompose (p a x) (q x b)

In category theory, existential quantification is encoded as the coend, which is a generalization of a colimit for profunctors. The coend formula for the composition of two profunctors reads:

(p ⊗ q) a b = ∫ z p a z × q z b

The product here is the cartesian product of sets.

Profunctors, being functors, form a category in which morphisms are natural transformations. As long as the two categories that they relate are the same, any two profunctors can be composed using a coend. So profunctor composition is a good candidate for a tensor product in such a category. It is indeed associative, up to isomorphism. But what’s the unit of profunctor composition? It turns out that the simplest profuctor — the hom-functor — because of the Yoneda lemma, is the unit of composition:

z C(a, z) × p z b ≅ p a b
∫ z p a z × C(z, b) ≅ p a b

Thus profunctors Cop×C->Set form a monoidal category.


Day Convolution

Or consider Set-valued functors. They can be composed using Day convolution. For that, the category C must itself be monoidal. Day convolution of two functors C->Set is defined using a coend:

(f ★ g) a = ∫ x y f x × g y × C(x ⊗ y, a)

Here, the tensor product of x ⊗ y comes from the monoidal category C, the other products are just cartesian products of sets (one of them being the hom-set).

As before, in Haskell, the coend turns into existential quantifier, which can be written symbolically:

Day f g a = exists x y. ((f x, g y), (x, y) -> a)

and encoded as a polymorphic constructor:

data Day f g a = forall x y. Day (f x) (g y) ((x, y) -> a)

We use the fact that the category of Haskell types is monoidal with respect to cartesian product.

We can build a monoidal category based on Day convolution. The unit with respect to Day convolution is C(i, -), the hom-functor applied to i — the unit in the monoidal category C. For instance, the left identity can be derived from:

(C(i, -) ★ g) a = ∫ x y C(i, x) × g y × C(x ⊗ y, a)

Applying the Yoneda lemma, or “integrating over x,” we get:

y g y × C(i ⊗ y, a)

Considering that i is the unit of the tensor product, we can perform the second integration to get g a.

The Monozoic Era

Monoidal categories are important because they provide rich grazing grounds for monoids. In a monoidal category we can define a more general monoid. It’s an object m with some special properties. These properties replace the usual definitions of multiplication and unit.

First, let’s reformulate the definition of a set-based monoid, taking into account the fact that Set is a monoidal category with respect to cartesian product.

A monoid is a set, so it’s an object in Set — let’s call it m. Multiplication maps pairs of elements of m back to m. These pairs are just elements of the cartesian product m × m. So multiplication is defined as a function:

μ :: m × m -> m

Unit of multiplication is a special element of m. We can select this element by providing a special morphism from the singleton set to m:

η :: () -> m

We can now express associativity and unit laws as properties of these two functions. The beauty of this formulation is that it generalizes easily to any cartesian category — just replace functions with morphisms and the unit () with the terminal object. There’s no reason to stop there: we can lift this definition all the way up to a monoidal category.

A monoid in a monoidal category is an object m together with two morphisms:

μ :: m ⊗ m -> m
η :: i -> m

Here i is the unit object with respect to the tensor product ⊗. Monoidal laws can be expressed using the associator α and the two unitors, λ and ρ, of the monoidal category:



Having previously defined several interesting monoidal categories, we can now go digging for new monoids.


Let’s start with the category of endofunctors where the tensor product is functor composition. A monoid in the category of endofunctors is an endofunctor m and two morphism. Remember that morphisms in a functor category are natural transformations. So we end up with two natural transformations:

μ :: m ∘ m -> m
η :: I -> m

where I is the identity functor. Their components at an object a are:

μa :: m (m a) -> m a
ηa :: a -> m a

This construct is easily recognizable as a monad. The associativity and unit laws are just monad laws. In Haskell, μa is called join and ηa is called return.


Let’s switch to the category of profunctors Cop×C->Set with profunctor composition as the tensor product. A monoid in that category is a profunctor ar. Multiplication is defined by a natural transformation:

μ :: ar ⊗ ar -> ar

Its component at a, b is:

μa b :: (∫ z ar a z × ar z b) -> ar a b

To simplify this formula we need a very useful identity that relates coends to ends. A hom-set that starts at a coend is equivalent to an end of the hom set:

C(∫ z p z z, y) ≅ ∫ z C(p z z, y)

Or, replacing external hom-sets with internal homs:

(∫ z p z z) -> y ≅ ∫ z (p z z -> y)

In Haskell, this formula is used to turn functions that take existential types to functions that are polymorphic:

(exists z. p z z) -> y ≅ forall z. (p z z -> y)

Intuitively, it makes perfect sense. If you want to define a function that takes an existential type, you have to be prepared to handle any type.

Using that identity, our multiplication formula can be rewritten as:

μa b :: ∫ z ((ar a z × ar z b) -> ar a b)

In Haskell, this derivation uses the existential quantifier:

mu a b = (exists z. (ar a z, ar z b)) -> ar a b

As we discussed, a function from an existential type is equivalent to a polymorphic function:

forall z. (ar a z, ar z b) -> ar a b

or, after currying and dropping the redundant quantification:

ar a z -> ar z b -> ar a b

This looks very much like a composition of morphisms in a category. In Haskell, this function is known in the infix-operator form as:

(>>>) :: ar a z -> ar z b -> ar a b

Let’s see what we get as the monoidal unit. Remember that the unit object in the profunctor category is the hom-functor C(a, b).

ηa b :: C(a, b) -> ar a b

In Haskell, this polymorphic function is traditionally called arr:

arr :: (a -> b) -> ar a b

The whole construct is known in Haskell as a pre-arrow. The full arrow is defined as a monoid in the category of strong profunctors, with strength defined as a natural transformation:

sta b :: p a b -> p (a, x) (b, x)

In Haskell, this function is called first.


There are several categorical formulations of what’s called in Haskell the applicative functor. To first approximaton, Haskell’s type system is the category Set. To translate Haskell constructs to category theory, the safest approach is to just play with endofunctors in Set. But both Set and its endofunctors have a lot of extra structure, so I’d like to start in a slightly more general setting.

Let’s have a look at the monoidal category of functors [C, Set], with Day convolution as the tensor product, and C(i, -) as unit. A monoid in this category is a functor f with multiplication given by the natural transformation:

μ :: f ★ f -> f

and unit given by:

η :: C(i, -) -> f

It turns out that the existence of these two natural transformations is equivalent to the requirement that f be a lax monoidal functor, which is the basis of the definition of the applicative functor in Haskell.

A monoidal functor is a functor that maps monoidal structure of one category to the monoidal structure of another category. It maps the tensor product, and it maps the unit object. In our case, the source category C has the monoidal structure given by the tensor product ⊗, and the target category Set is monoidal with respect to the cartesian product ×. A functor is monoidal if it doesn’t matter whether we first map two object and then multiply them, or first multiply them and then map the result:

f x × f y ≅ f (x ⊗ y)

Also, the unit object in Set should be isomporphic to the result of mapping the unit object in C:

() ≅ f i

Here, () is the terminal object in Set and i is the unit object in C.

These conditions are relaxed in the definition of a lax monoidal functor. A lax monoidal functor replaces isomorphisms with regular unidirectional morphisms:

f x × f y -> f (x ⊗ y)
() -> f i

It can be shown that the monoid in the category[C, Set], with Day convolution as the tensor product, is equivalent to the lax monoidal functor.

The Haskell definition of Applicative doesn’t look like Day convolution or like a lax monoidal functor:

class Functor f => Applicative f where
    (<*>) :: f (a -> b) -> (f a -> f b)
    pure :: a -> f a

You may recognize pure as a component of η, the natural transformation defining the monoid with respect to Day convolution. When you replace the category C with Set, the unit object C(i, -) turns into the identity functor. However, the operator <*> is lifted from the definition of yet another lax functor, the lax closed functor. It’s a functor that preserves the closed structure defined by the internal hom functor. In Set, the internal hom functor is just the arrow (->), hence the definition:

class Functor f => Closed f where
    (<*>) :: f (a -> b) -> (f a -> f b)
    unit :: f ()

As long as the internal hom is defined through the adjunction with the product, a lax closed functor is equivalent to a lax monoidal functor.


It is pretty shocking to realize how many different animals share the same body plan — I’m talking here about the monoid as the skeleton of a myriad of different mathematical and programming constructs. And I haven’t even touched on the whole kingdom of enriched categories, where monoidal categories form the reservoir of hom-objects. Virtually all notions I’ve discussed here can be generalized to enriched categories, including functors, profunctors, the Yoneda lemma, Day convolution, and so on.


  • Hadean Eon: Began with the formation of the Earth about 4.6 billion years ago. It’s the period before the earliest-known rocks.
  • Archean Eon: During the Archean, the Earth’s crust had cooled enough to allow the formation of continents.
  • Cambrian explosion: Relatively short evolutionary event, during which most major animal phyla appeared.
  • Arthropods: from Greek ἄρθρωσις árthrosis, “joint”
  • Tensor, from Latin tendere “to stretch”
  • Functor: from Latin fungi, “perform”

Unlike monads, which came into programming straight from category theory, applicative functors have their origins in programming. McBride and Paterson introduced applicative functors as a programming pearl in their paper Applicative programming with effects. They also provided a categorical interpretation of applicatives in terms of strong lax monoidal functors. It’s been accepted that, just like “a monad is a monoid in the category of endofunctors,” so “an applicative is a strong lax monoidal functor.”

The so called “tensorial strength” seems to be important in categorical semantics, and in his seminal paper Notions of computation and monads, Moggi argued that effects should be described using strong monads. It makes sense, considering that a computation is done in a context, and you should be able to make the global context available under the monad. The fact that we don’t talk much about strong monads in Haskell is due to the fact that all functors in the category Set, which underlies the Haskell’s type system, have canonical strength. So why do we talk about strength when dealing with applicative functors? I have looked into this question and came to the conclusion that there is no fundamental reason, and that it’s okay to just say:

An applicative is a lax monoidal functor

In this post I’ll discuss different equivalent categorical definitions of the applicative functor. I’ll start with a lax closed functor, then move to a lax monoidal functor, and show the equivalence of the two definitions. Then I’ll introduce the calculus of ends and show that the third definition of the applicative functor as a monoid in a suitable functor category equipped with Day convolution is equivalent to the previous ones.

Applicative as a Lax Closed Functor

The standard definition of the applicative functor in Haskell reads:

class Functor f => Applicative f where
    (<*>) :: f (a -> b) -> (f a -> f b)
    pure :: a -> f a

At first sight it doesn’t seem to involve a monoidal structure. It looks more like preserving function arrows (I added some redundant parentheses to suggest this interpretation).

Categorically, functors that “preserve arrows” are known as closed functors. Let’s look at a definition of a closed functor f between two categories C and D. We have to assume that both categories are closed, meaning that they have internal hom-objects for every pair of objects. Internal hom-objects are also called function objects or exponentials. They are normally defined through the right adjoint to the product functor:

C(z × a, b) ≅ C(z, a => b)

To distinguish between sets of morphisms and function objects (they are the same thing in Set), I will temporarily use double arrows for function objects.

We can take a functor f and act with it on the function object a=>b in the category C. We get an object f (a=>b) in D. Or we can map the two objects a and b from C to D and then construct the function object in D: f a => f b.


We call a functor closed if the two results are isomorphic (I have subscripted the two arrows with the categories where they are defined):

f (a =>C b) ≅ (f a =>D f b)

and if the functor preserves the unit object:

iD ≅ f iC

What’s the unit object? Normally, this is the unit with respect to the same product that was used to define the function object using the adjunction. I’m saying “normally,” because it’s possible to define a closed category without a product.

Note: The two arrows and the two is are defined with respect to two different products. The first isomorphism must be natural in both a and b. Also, to complete the picture, there are some diagrams that must commute.

The two isomorphisms that define a closed functor can be relaxed and replaced by unidirectional morphisms. The result is a lax closed functor:

f (a => b) -> (f a => f b)
i -> f i

This looks almost like the definition of Applicative, except for one problem: how can we recover the natural transformation we call pure from a single morphism i -> f i.

One way to do it is from the position of strength. An endofunctor f has tensorial strength if there is a natural transformation:

stc a :: c ⊗ f a -> f (c ⊗ a)

Think of c as the context in which the computation f a is performed. Strength means that we can use this external context inside the computation.

In the category Set, with the tensor product replaced by cartesian product, all functors have canonical strength. In Haskell, we would define it as:

st (c, fa) = fmap ((,) c) fa

The morphism in the definition of the lax closed functor translates to:

unit :: () -> f ()

Notice that, up to isomorphism, the unit type () is the unit with respect to cartesian product. The relevant isomorphisms are:

λa :: ((), a) -> a
ρa :: (a, ()) -> a

Here’s the derivation from Rivas and Jaskelioff’s Notions of Computation as Monoids:

≅  (a, ())   -- unit law, ρ-1
-> (a, f ()) -- lax unit
-> f (a, ()) -- strength
≅  f a       -- lifted unit law, f ρ

Strength is necessary if you’re starting with a lax closed (or monoidal — see the next section) endofunctor in an arbitrary closed (or monoidal) category and you want to derive pure within that category — not after you restrict it to Set.

There is, however, an alternative derivation using the Yoneda lemma:

f ()
≅ forall a. (() -> a) -> f a  -- Yoneda
≅ forall a. a -> f a -- because: (() -> a) ≅ a

We recover the whole natural transformation from a single value. The advantage of this derivation is that it generalizes beyond endofunctors and it doesn’t require strength. As we’ll see later, it also ties nicely with the Day-convolution definition of applicative. The Yoneda lemma only works for Set-valued functors, but so does Day convolution (there are enriched versions of both Yoneda and Day convolution, but I’m not going to discuss them here).

We can define the categorical version of the Haskell’s applicative functor as a lax closed functor going from a closed category C to Set. It’s a functor equipped with a natural transformation:

f (a => b) -> (f a -> f b)

where a=>b is the internal hom-object in C (the second arrow is a function type in Set), and a function:

1 -> f i

where 1 is the singleton set and i is the unit object in C.

The importance of a categorical definition is that it comes with additional identities or “axioms.” A lax closed functor must be compatible with the structure of both categories. I will not go into details here, because we are really only interested in closed categories that are monoidal, where these axioms are easier to express.

The definition of a lax closed functor is easily translated to Haskell:

class Functor f => Closed f where
    (<*>) :: f (a -> b) -> f a -> f b
    unit :: f ()

Applicative as a Lax Monoidal Functor

Even though it’s possible to define a closed category without a monoidal structure, in practice we usually work with monoidal categories. This is reflected in the equivalent definition of the Haskell’s applicative functor as a lax monoidal functor. In Haskell, we would write:

class Functor f => Monoidal f where
    (>*<) :: (f a, f b) -> f (a, b)
    unit :: f ()

This definition is equivalent to our previous definition of a closed functor. That’s because, as we’ve seen, a function object in a monoidal category is defined in terms of a product. We can show the equivalence in a more general categorical setting.

This time let’s start with a symmetric closed monoidal category C, in which the function object is defined through the right adjoint to the tensor product:

C(z ⊗ a, b) ≅ C(z, a => b)

As usual, the tensor product is associative and unital — with the unit object i — up to isomorphism. The symmetry is defined through natural isomorphism:

γ :: a ⊗ b -> b ⊗ a

A functor f between two monoidal categories is lax monoidal if there exist: (1) a natural transformation

f a ⊗ f b -> f (a ⊗ b)

and (2) a morphism

i -> f i

Notice that the products and units on either side of the two mappings are from different categories.

A (lax-) monoidal functor must also preserve associativity and unit laws.

For instance a triple product

f a ⊗ (f b ⊗ f c)

may be rearranged using an associator α to give

(f a ⊗ f b) ⊗ f c

then converted to

f (a ⊗ b) ⊗ f c

and then to

f ((a ⊗ b) ⊗ c)

Or it could be first converted to

f a ⊗ f (b ⊗ c)

and then to

f (a ⊗ (b ⊗ c))

These two should be equivalent under the associator in C.


Similarly, f a ⊗ i can be simplified to f a using the right unitor ρ in D. Or it could be first converted to f a ⊗ f i, then to f (a ⊗ i), and then to f a, using the right unitor in C. The two paths should be equivalent. (Similarly for the left identity.)


We will now consider functors from C to Set, with Set equipped with the usual cartesian product, and the singleton set as unit. A lax monoidal functor is defined by: (1) a natural transformation:

(f a, f b) -> f (a ⊗ b)

and (2) a choice of an element of the set f i (a function from 1 to f i picks an element from that set).

We need the target category to be Set because we want to be able to use the Yoneda lemma to show equivalence with the standard definition of applicative. I’ll come back to this point later.

The Equivalence

The definitions of a lax closed and a lax monoidal functors are equivalent when C is a closed symmetric monoidal category. The proof relies on the existence of the adjunction, in particular the unit and the counit of the adjunction:

ηa :: a -> (b => (a ⊗ b))
εb :: (a => b) ⊗ a -> b

For instance, let’s assume that f is lax-closed. We want to construct the mapping

(f a, f b) -> f (a ⊗ b)

First, we apply the lifted pair (unit, identity), (f η, f id)

(f a -> f (b => a ⊗ b), f id)

to the left hand side. We get:

(f (b => a ⊗ b), f b)

Now we can use (the uncurried version of) the lax-closed morphism:

(f (b => x), f b) -> f x

to get:

f (a ⊗ b)

Conversely, assuming the lax-monoidal property we can show that the functor is lax-closed, that is to say, implement the following function:

(f (a => b), f a) -> f b

First we use the lax monoidal morphism on the left hand side:

f ((a => b) ⊗ a)

and then use the counit (a.k.a. the evaluation morphism) to get the desired result f b

There is yet another presentation of applicatives using Day convolution. But before we get there, we need a little refresher on calculus.

Calculus of Ends

Ends and coends are very useful constructs generalizing limits and colimits. They are defined through universal constructions. They have a few fundamental properties that are used over and over in categorical calculations. I’ll just introduce the notation and a few important identities. We’ll be working in a symmetric monoidal category C with functors from C to Set and profunctors from Cop×C to Set. The end of a profunctor p is a set denoted by:

a p a a

The most important thing about ends is that a set of natural transformations between two functors f and g can be represented as an end:

[C, Set](f, g) = ∫a C(f a, g a)

In Haskell, the end corresponds to universal quantification over a functor of mixed variance. For instance, the natural transformation formula takes the familiar form:

forall a. f a -> g a

The Yoneda lemma, which deals with natural transformations, can also be written using an end:

z (C(a, z) -> f z) ≅ f a

In Haskell, we can write it as the equivalence:

forall z. ((a -> z) -> f z) ≅ f a

which is a generalization of the continuation passing transform.

The dual notion of coend is similarly written using an integral sign, with the “integration variable” in the superscript position:

a p a a

In pseudo-Haskell, a coend is represented by an existential quantifier. It’s possible to define existential data types in Haskell by converting existential quantification to universal one. The relevant identity in terms of coends and ends reads:

(∫ z p z z) -> y ≅ ∫ z (p z z -> y)

In Haskell, this formula is used to turn functions that take existential types to functions that are polymorphic:

(exists z. p z z) -> y ≅ forall z. (p z z -> y)

Intuitively, it makes perfect sense. If you want to define a function that takes an existential type, you have to be prepared to handle any type.

The equivalent of the Yoneda lemma for coends reads:

z f z × C(z, a) ≅ f a

or, in pseudo-Haskell:

exists z. (f z, z -> a) ≅ f a

(The intuition is that the only thing you can do with this pair is to fmap the function over the first component.)

There is also a contravariant version of this identity:

z C(a, z) × f z ≅ f a

where f is a contravariant functor (a.k.a., a presheaf). In pseudo-Haskell:

exists z. (a -> z, f z) ≅ f a

(The intuition is that the only thing you can do with this pair is to apply the contramap of the first component to the second component.)

Using coends we can define a tensor product in the category of functors [C, Set]. This product is called Day convolution:

(f ★ g) a = ∫ x y f x × g y × C(x ⊗ y, a)

It is a bifunctor in that category (read, it can be used to lift natural transformations). It’s associative and symmetric up to isomorphism. It also has a unit — the hom-functor C(i, -), where i is the monoidal unit in C. In other words, Day convolution imbues the category [C, Set] with monoidal structure.

Let’s verify the unit laws.

(C(i, -) ★ g) a = ∫ x y C(i, x) × g y × C(x ⊗ y, a)

We can use the contravariant Yoneda to “integrate over x” to get:

y g y × C(i ⊗ y, a)

Considering that i is the unit of the tensor product in C, we get:

y g y × C(y, a)

Covariant Yoneda lets us “integrate over y” to get the desired g a. The same method works for the right unit law.

Applicative as a Monoid

Given a monoidal category, we can always define a monoid as an object m equipped with two morphisms:

μ :: m ⊗ m -> m
η :: i -> m

satisfying the laws of associativity and unitality.

We have shown that the functor category [C, Set] (with C a symmetric monoidal category) is monoidal under Day convolution. An object in this category is a functor f. The two morphisms that would make it a candidate for a monoid are natural transformations:

μ :: f ★ f -> f
η :: C(i, -) -> f

The a component of the natural transformation μ can be rewritten as:

(∫ x y f x × f y × C(x ⊗ y, a)) -> f a

which is equivalent to:

x y (f x × f y × C(x ⊗ y, a) -> f a)

or, upon currying:

x y (f x, f y) -> C(x ⊗ y, a) -> f a

It turns out that so defined monoid is equivalent to a lax monoidal functor. This was shown by Rivas and Jaskelioff. The following derivation is due to Bob Atkey.

The trick is to start with the whole set of natural transformation from f★f to f. The multiplication μ is just one of them. We’ll express the set of natural transformations as an end:

a ((f ★ f) a -> f a)

Plugging in the formula for the a component of μ, we get:

a x y (f x, f y) -> C(x ⊗ y, a) -> f a

The end over a does not involve the first argument, so we can move the integral sign:

x y (f x, f y) -> ∫ a C(x ⊗ y, a) -> f a

Then we use the Yoneda lemma to “perform the integration” over a:

x y (f x, f y) -> f (x ⊗ y)

You may recognize this as a set of natural transformations that define a lax monoidal functor. We have established a one-to-one correspondence between these natural transformations and the ones defining monoidal multiplication using Day convolution.

The remaining part is to show the equivalence between the unit with respect to Day convolution and the second part of the definition of the lax monoidal functor, the morphism:

1 -> f i

We start with the set of natural transformations that contains our η:

a (i -> a) -> f a

By Yoneda, this is just f i. Picking an element from a set is equivalent to defining a morphism from the singleton set 1, so for any choice of η we get:

1 -> f i

and vice versa. The two definitions are equivalent.

Notice that the monoidal unit η under Day convolution becomes the definition of pure in the Haskell version of applicative. Indeed, when we replace the category C with Set, f becomes and endofunctor, and the unit of Day convolution C(i, -) becomes the identity functor Id. We get:

η :: Id -> f

or, in components:

pure :: a -> f a

So, strictly speaking, the Haskell definition of Applicative mixes the elements of the lax closed functor and the monoidal unit under Day convolution.


I’m grateful to Mauro Jaskelioff and Exequiel Rivas for correspondence and to Bob Atkey, Dimitri Chikhladze, and Make Shulman for answering my questions on Math Overflow.

This is part 23 of Categories for Programmers. Previously: Monads Categorically. See the Table of Contents.

Now that we have covered monads, we can reap the benefits of duality and get comonads for free simply by reversing the arrows and working in the opposite category.

Recall that, at the most basic level, monads are about composing Kleisli arrows:

a -> m b

where m is a functor that is a monad. If we use the letter w (upside down m) for the comonad, we can define co-Kleisli arrows as morphism of the type:

w a -> b

The analog of the fish operator for co-Kleisli arrows is defined as:

(=>=) :: (w a -> b) -> (w b -> c) -> (w a -> c)

For co-Kleisli arrows to form a category we also have to have an identity co-Kleisli arrow, which is called extract:

extract :: w a -> a

This is the dual of return. We also have to impose the laws of associativity as well as left- and right-identity. Putting it all together, we could define a comonad in Haskell as:

class Functor w => Comonad w where
    (=>=) :: (w a -> b) -> (w b -> c) -> (w a -> c)
    extract :: w a -> a

In practice, we use slightly different primitives, as we’ll see shortly.

The question is, what’s the use for comonads in programming?

Programming with Comonads

Let’s compare the monad with the comonad. A monad provides a way of putting a value in a container using return. It doesn’t give you access to a value or values stored inside. Of course, data structures that implement monads might provide access to their contents, but that’s considered a bonus. There is no common interface for extracting values from a monad. And we’ve seen the example of the IO monad that prides itself in never exposing its contents.

A comonad, on the other hand, provides the means of extracting a single value from it. It does not give the means to insert values. So if you want to think of a comonad as a container, it always comes pre-filled with contents, and it lets you peek at it.

Just as a Kleisli arrow takes a value and produces some embellished result — it embellishes it with context — a co-Kleisli arrow takes a value together with a whole context and produces a result. It’s an embodiment of contextual computation.

The Product Comonad

Remember the reader monad? We introduced it to tackle the problem of implementing computations that need access to some read-only environment e. Such computations can be represented as pure functions of the form:

(a, e) -> b

We used currying to turn them into Kleisli arrows:

a -> (e -> b)

But notice that these functions already have the form of co-Kleisli arrows. Let’s massage their arguments into the more convenient functor form:

data Product e a = P e a
  deriving Functor

We can easily define the composition operator by making the same environment available to the arrows that we are composing:

(=>=) :: (Product e a -> b) -> (Product e b -> c) -> (Product e a -> c)
f =>= g = \(P e a) -> let b = f (P e a)
                          c = g (P e b)
                       in c

The implementation of extract simply ignores the environment:

extract (P e a) = a

Not surprisingly, the product comonad can be used to perform exactly the same computations as the reader monad. In a way, the comonadic implementation of the environment is more natural — it follows the spirit of “computation in context.” On the other hand, monads come with the convenient syntactic sugar of the do notation.

The connection between the reader monad and the product comonad goes deeper, having to do with the fact that the reader functor is the right adjoint of the product functor. In general, though, comonads cover different notions of computation than monads. We’ll see more examples later.

It’s easy to generalize the Product comonad to arbitrary product types including tuples and records.

Dissecting the Composition

Continuing the process of dualization, we could go ahead and dualize monadic bind and join. Alternatively, we can repeat the process we used with monads, where we studied the anatomy of the fish operator. This approach seems more enlightening.

The starting point is the realization that the composition operator must produce a co-Kleisli arrow that takes w a and produces a c. The only way to produce a c is to apply the second function to an argument of the type w b:

(=>=) :: (w a -> b) -> (w b -> c) -> (w a -> c)
f =>= g = g ... 

But how can we produce a value of type w b that could be fed to g? We have at our disposal the argument of type w a and the function f :: w a -> b. The solution is to define the dual of bind, which is called extend:

extend :: (w a -> b) -> w a -> w b

Using extend we can implement composition:

f =>= g = g . extend f

Can we next dissect extend? You might be tempted to say, why not just apply the function w a -> b to the argument w a, but then you quickly realize that you’d have no way of converting the resulting b to w b. Remember, the comonad provides no means of lifting values. At this point, in the analogous construction for monads, we used fmap. The only way we could use fmap here would be if we had something of the type w (w a) at our disposal. If we coud only turn w a into w (w a). And, conveniently, that would be exactly the dual of join. We call it duplicate:

duplicate :: w a -> w (w a)

So, just like with the definitions of the monad, we have three equivalent definitions of the comonad: using co-Kleisli arrows, extend, or duplicate. Here’s the Haskell definition taken directly from Control.Comonad library:

class Functor w => Comonad w where
  extract :: w a -> a
  duplicate :: w a -> w (w a)
  duplicate = extend id
  extend :: (w a -> b) -> w a -> w b
  extend f = fmap f . duplicate

Provided are the default implementations of extend in terms of duplicate and vice versa, so you only need to override one of them.

The intuition behind these functions is based on the idea that, in general, a comonad can be thought of as a container filled with values of type a (the product comonad was a special case of just one value). There is a notion of the “current” value, one that’s easily accessible through extract. A co-Kleisli arrow performs some computation that is focused on the current value, but it has access to all the surrounding values. Think of the Conway’s game of life. Each cell contains a value (usually just True or False). A comonad corresponding to the game of life would be a grid of cells focused on the “current” cell.

So what does duplicate do? It takes a comonadic container w a and produces a container of containers w (w a). The idea is that each of these containers is focused on a different a inside w a. In the game of life, you would get a grid of grids, each cell of the outer grid containing an inner grid that’s focused on a different cell.

Now look at extend. It takes a co-Kleisli arrow and a comonadic container w a filled with as. It applies the computation to all of these as, replacing them with bs. The result is a comonadic container filled with bs. extend does it by shifting the focus from one a to another and applying the co-Kleisli arrow to each of them in turn. In the game of life, the co-Kleisli arrow would calculate the new state of the current cell. To do that, it would look at its context — presumably its nearest neighbors. The default implementation of extend illustrates this process. First we call duplicate to produce all possible foci and then we apply f to each of them.

The Stream Comonad

This process of shifting the focus from one element of the container to another is best illustrated with the example of an infinite stream. Such a stream is just like a list, except that it doesn’t have the empty constructor:

data Stream a = Cons a (Stream a)

It’s trivially a Functor:

instance Functor Stream where
    fmap f (Cons a as) = Cons (f a) (fmap f as)

The focus of a stream is its first element, so here’s the implementation of extract:

extract (Cons a _) = a

duplicate produces a stream of streams, each focused on a different element.

duplicate (Cons a as) = Cons (Cons a as) (duplicate as)

The first element is the original stream, the second element is the tail of the original stream, the third element is its tail, and so on, ad infinitum.

Here’s the complete instance:

instance Comonad Stream where
    extract (Cons a _) = a
    duplicate (Cons a as) = Cons (Cons a as) (duplicate as)

This is a very functional way of looking at streams. In an imperative language, we would probably start with a method advance that shifts the stream by one position. Here, duplicate produces all shifted streams in one fell swoop. Haskell’s laziness makes this possible and even desirable. Of course, to make a Stream practical, we would also implement the analog of advance:

tail :: Stream a -> Stream a
tail (Cons a as) = as

but it’s never part of the comonadic interface.

If you had any experience with digital signal processing, you’ll see immediately that a co-Kleisli arrow for a stream is just a digital filter, and extend produces a filtered stream.

As a simple example, let’s implement the moving average filter. Here’s a function that sums n elements of a stream:

sumS :: Num a => Int -> Stream a -> a
sumS n (Cons a as) = if n <= 0 then 0 else a + sumS (n - 1) as

Here’s the function that calculates the average of the first n elements of the stream:

average :: Fractional a => Int -> Stream a -> a
average n stm = (sumS n stm) / (fromIntegral n)

Partially applied average n is a co-Kleisli arrow, so we can extend it over the whole stream:

movingAvg :: Fractional a => Int -> Stream a -> Stream a
movingAvg n = extend (average n)

The result is the stream of running averages.

A stream is an example of a unidirectional, one-dimensional comonad. It can be easily made bidirectional or extended to two or more dimensions.

Comonad Categorically

Defining a comonad in category theory is a straightforward exercise in duality. As with the monad, we start with an endofunctor T. The two natural transformations, η and μ, that define the monad are simply reversed for the comonad:

ε :: T -> I
δ :: T -> T2

The components of these transformations correspond to extract and duplicate. Comonad laws are the mirror image of monad laws. No big surprise here.

Then there is the derivation of the monad from an adjunction. Duality reverses an adjunction: the left adjoint becomes the right adjoint and vice versa. And, since the composition R ∘ L defines a monad, L ∘ R must define a comonad. The counit of the adjunction:

ε :: L ∘ R -> I

is indeed the same ε that we see in the definition of the comonad — or, in components, as Haskell’s extract. We can also use the unit of the adjunction:

η :: I -> R ∘ L

to insert an R ∘ L in the middle of L ∘ R and produce L ∘ R ∘ L ∘ R. Making T2 from T defines the δ, and that completes the definition of the comonad.

We’ve also seen that the monad is a monoid. The dual of this statement would require the use of a comonoid, so what’s a comonoid? The original definition of a monoid as a single-object category doesn’t dualize to anything interesting. When you reverse the direction of all endomorphisms, you get another monoid. Recall, however, that in our approach to a monad, we used a more general definition of a monoid as an object in a monoidal category. The construction was based on two morphisms:

μ :: m ⊗ m -> m
η :: i -> m

The reversal of these morphisms produces a comonoid in a monoidal category:

δ :: m -> m ⊗ m
ε :: m -> i

One can write a definition of a comonoid in Haskell:

class Comonoid m where
  split   :: m -> (m, m)
  destroy :: m -> ()

but it is rather trivial. Obviously destroy ignores its argument.

destroy _ = ()

split is just a pair of functions:

split x = (f x, g x)

Now consider comonoid laws that are dual to the monoid unit laws.

lambda . bimap destroy id . split = id
rho . bimap id destroy . split = id

Here, lambda and rho are the left and right unitors, respectively (see the definition of monoidal categories). Plugging in the definitions, we get:

lambda (bimap destroy id (split x))
= lambda (bimap destroy id (f x, g x))
= lambda ((), g x)
= g x

which proves that g = id. Similarly, the second law expands to f = id. In conclusion:

split x = (x, x)

which shows that in Haskell (and, in general, in the category Set) every object is a trivial comonoid.

Fortunately there are other more interesting monoidal categories in which to define comonoids. One of them is the category of endofunctors. And it turns out that, just like the monad is a monoid in the category of endofunctors,

The comonad is a comonoid in the category of endofunctors.

The Store Comonad

Another important example of a comonad is the dual of the state monad. It’s called the costate comonad or, alternatively, the store comonad.

We’ve seen before that the state monad is generated by the adjunction that defines the exponentials:

L z = z × s
R a = s ⇒ a

We’ll use the same adjunction to define the costate comonad. A comonad is defined by the composition L ∘ R:

L (R a) = (s ⇒ a) × s

Translating this to Haskell, we start with the adjunction between the Prod functor on the left and the Reader functor or the right. Composing Prod after Reader is equivalent to the following definition:

data Store s a = Store (s -> a) s

The counit of the adjunction taken at the object a is the morphism:

εa :: ((s ⇒ a) × s) -> a

or, in Haskell notation:

counit (Prod (Reader f, s)) = f s

This becomes our extract:

extract (Store f s) = f s

The unit of the adjunction:

unit a = Reader (\s -> Prod (a, s))

can be rewritten as partially applied data constructor:

Store f :: s -> Store f s

We construct δ, or duplicate, as the horizontal composition:

δ :: L ∘ R -> L ∘ R ∘ L ∘ R
δ = L ∘ η ∘ R

We have to sneak η through the leftmost L, which is the Prod functor. It means acting with η, or Store f, on the left component of the pair (that’s what fmap for Prod would do). We get:

duplicate (Store f s) = Store (Store f) s

(Remember that, in the formula for δ, L and R stand for identity natural transformations whose components are identity morphisms.)

Here’s the complete definition of the Store comonad:

instance Comonad (Store s) where
  extract (Store f s) = f s
  duplicate (Store f s) = Store (Store f) s

You may think of the Reader part of Store as a generalized container of as that are keyed using elements of the type s. For instance, if s is Int, Reader Int a is an infinite bidirectional stream of as. Store pairs this container with a value of the key type. For instance, Reader Int a is paired with an Int. In this case, extract uses this integer to index into the infinite stream. You may think of the second component of Store as the current position.

Continuing with this example, duplicate creates a new infinite stream indexed by an Int. This stream contains streams as its elements. In particular, at the current position, it contains the original stream. But if you use some other Int (positive or negative) as the key, you’d obtain a shifted stream positioned at that new index.

In general, you can convince yourself that when extract acts on the duplicated Store it produces the original Store (in fact, the identity law for the comonad states that extract . duplicate = id).

The Store comonad plays an important role as the theoretical basis for the Lens library. Conceptually, the Store s a comonad encapsulates the idea of “focusing” (like a lens) on a particular substructure of the date type a using the type s as an index. In particular, a function of the type:

a -> Store s a

is equivalent to a pair of functions:

set :: a -> s -> a
get :: a -> s

If a is a product type, set could be implemented as setting the field of type s inside of a while returning the modified version of a. Similarly, get could be implemented to read the value of the s field from a. We’ll explore these ideas more in the next section.


  1. Implement the Conway’s Game of Life using the Store comonad. Hint: What type do you pick for s?


I’m grateful to Edward Kmett for reading the draft of this post and pointing out flaws in my reasoning.

Next: F-Algebras.

This is part 20 of Categories for Programmers. Previously: Free/Forgetful Adjunctions. See the Table of Contents.

Programmers have developed a whole mythology around monads. It’s supposed to be one of the most abstract and difficult concepts in programming. There are people who “get it” and those who don’t. For many, the moment when they understand the concept of the monad is like a mystical experience. The monad abstracts the essence of so many diverse constructions that we simply don’t have a good analogy for it in everyday life. We are reduced to groping in the dark, like those blind men touching different parts of the elephant end exclaiming triumphantly: “It’s a rope,” “It’s a tree trunk,” or “It’s a burrito!”

Let me set the record straight: The whole mysticism around the monad is the result of a misunderstanding. The monad is a very simple concept. It’s the diversity of applications of the monad that causes the confusion.

As part of research for this post I looked up duct tape (a.k.a., duck tape) and its applications. Here’s a little sample of things that you can do with it:

  • sealing ducts
  • fixing CO2 scrubbers on board Apollo 13
  • wart treatment
  • fixing Apple’s iPhone 4 dropped call issue
  • making a prom dress
  • building a suspension bridge

Now imagine that you didn’t know what duct tape was and you were trying to figure it out based on this list. Good luck!

So I’d like to add one more item to the collection of “the monad is like…” clichés: The monad is like duct tape. Its applications are widely diverse, but its principle is very simple: it glues things together. More precisely, it composes things.

This partially explains the difficulties a lot of programmers, especially those coming from the imperative background, have with understanding the monad. The problem is that we are not used to thinking of programing in terms of function composition. This is understandable. We often give names to intermediate values rather than pass them directly from function to function. We also inline short segments of glue code rather than abstract them into helper functions. Here’s an imperative-style implementation of the vector-length function in C:

double vlen(double * v) {
  double d = 0.0;
  int n;
  for (n = 0; n < 3; ++n)
    d += v[n] * v[n];
  return sqrt(d);

Compare this with the (stylized) Haskell version that makes function composition explicit:

vlen = sqrt . sum . fmap  (flip (^) 2)

(Here, to make things even more cryptic, I partially applied the exponentiation operator (^) by setting its second argument to 2.)

I’m not arguing that Haskell’s point-free style is always better, just that function composition is at the bottom of everything we do in programming. And even though we are effectively composing functions, Haskell does go to great lengths to provide imperative-style syntax called the do notation for monadic composition. We’ll see its use later. But first, let me explain why we need monadic composition in the first place.

The Kleisli Category

We have previously arrived at the writer monad by embellishing regular functions. The particular embellishment was done by pairing their return values with strings or, more generally, with elements of a monoid. We can now recognize that such embellishment is a functor:

newtype Writer w a = Writer (a, w)

instance Functor (Writer w) where
  fmap f (Writer (a, w)) = Writer (f a, w)

We have subsequently found a way of composing embellished functions, or Kleisli arrows, which are functions of the form:

a -> Writer w b

It was inside the composition that we implemented the accumulation of the log.

We are now ready for a more general definition of the Kleisli category. We start with a category C and an endofunctor m. The corresponding Kleisli category K has the same objects as C, but its morphisms are different. A morphism between two objects a and b in K is implemented as a morphism:

a -> m b

in the original category C. It’s important to keep in mind that we treat a Kleisli arrow in K as a morphism between a and b, and not between a and m b.

In our example, m was specialized to Writer w, for some fixed monoid w.

Kleisli arrows form a category only if we can define proper composition for them. If there is a composition, which is associative and has an identity arrow for every object, then the functor m is called a monad, and the resulting category is called the Kleisli category.

In Haskell, Kleisli composition is defined using the fish operator >=>, and the identity arrrow is a polymorphic function called return. Here’s the definition of a monad using Kleisli composition:

class Monad m where
  (>=>) :: (a -> m b) -> (b -> m c) -> (a -> m c)
  return :: a -> m a

Keep in mind that there are many equivalent ways of defining a monad, and that this is not the primary one in the Haskell ecosystem. I like it for its conceptual simplicity and the intuition it provides, but there are other definitions that are more convenient when programming. We’ll talk about them momentarily.

In this formulation, monad laws are very easy to express. They cannot be enforced in Haskell, but they can be used for equational reasoning. They are simply the standard composition laws for the Kleisli category:

(f >=> g) >=> h = f >=> (g >=> h) -- associativity
return >=> f = f                  -- left unit
f >=> return = f                  -- right unit

This kind of a definition also expresses what a monad really is: it’s a way of composing embellished functions. It’s not about side effects or state. It’s about composition. As we’ll see later, embellished functions may be used to express a variety of effects or state, but that’s not what the monad is for. The monad is the sticky duct tape that ties one end of an embellished function to the other end of an embellished function.

Going back to our Writer example: The logging functions (the Kleisli arrows for the Writer functor) form a category because Writer is a monad:

instance Monoid w => Monad (Writer w) where
    f >=> g = \a -> 
        let Writer (b, s)  = f a
            Writer (c, s') = g b
        in Writer (c, s `mappend` s')
    return a = Writer (a, mempty)

Monad laws for Writer w are satisfied as long as monoid laws for w are satisfied (they can’t be enforced in Haskell either).

There’s a useful Kleisli arrow defined for the Writer monad called tell. It’s sole purpose is to add its argument to the log:

tell :: w -> Writer w ()
tell s = Writer ((), s)

We’ll use it later as a building block for other monadic functions.

Fish Anatomy

When implementing the fish operator for different monads you quickly realize that a lot of code is repeated and can be easily factored out. To begin with, the Kleisli composition of two functions must return a function, so its implementation may as well start with a lambda taking an argument of type a:

(>=>) :: (a -> m b) -> (b -> m c) -> (a -> m c)
f >=> g = \a -> ...

The only thing we can do with this argument is to pass it to f:

f >=> g = \a -> let mb = f a
                in ...

At this point we have to produce the result of type m c, having at our disposal an object of type m b and a function g :: b -> m c. Let’s define a function that does that for us. This function is called bind and is usually written in the form of an infix operator:

(>>=) :: m a -> (a -> m b) -> m b

For every monad, instead of defining the fish operator, we may instead define bind. In fact the standard Haskell definition of a monad uses bind:

class Monad m where
    (>>=) :: m a -> (a -> m b) -> m b
    return :: a -> m a

Here’s the definition of bind for the Writer monad:

(Writer (a, w)) >>= f = let Writer (b, w') = f a
                        in  Writer (b, w `mappend` w')

It is indeed shorter than the definition of the fish operator.

It’s possible to further dissect bind, taking advantage of the fact that m is a functor. We can use fmap to apply the function a -> m b to the contents of m a. This will turn a into m b. The result of the application is therefore of type m (m b). This is not exactly what we want — we need the result of type m b — but we’re close. All we need is a function that collapses or flattens the double application of m. Such function is called join:

join :: m (m a) -> m a

Using join, we can rewrite bind as:

ma >>= f = join (fmap f ma)

That leads us to the third option for defining a monad:

class Functor m => Monad m where
    join :: m (m a) -> m a
    return :: a -> m a

Here we have explicitly requested that m be a Functor. We didn’t have to do that in the previous two definitions of the monad. That’s because any type constructor m that either supports the fish or bind operator is automatically a functor. For instance, it’s possible to define fmap in terms of bind and return:

fmap f ma = ma >>= \a -> return (f a)

For completeness, here’s join for the Writer monad:

join :: Monoid w => Writer w (Writer w a) -> Writer w a
join (Writer ((Writer (a, w')), w)) = Writer (a, w `mappend` w')

The do Notation

One way of writing code using monads is to work with Kleisli arrows — composing them using the fish operator. This mode of programming is the generalization of the point-free style. Point-free code is compact and often quite elegant. In general, though, it can be hard to understand, bordering on cryptic. That’s why most programmers prefer to give names to function arguments and intermediate values.

When dealing with monads it means favoring the bind operator over the fish operator. Bind takes a monadic value and returns a monadic value. The programmer may chose to give names to those values. But that’s hardly an improvement. What we really want is to pretend that we are dealing with regular values, not the monadic containers that encapsulate them. That’s how imperative code works — side effects, such as updating a global log, are mostly hidden from view. And that’s what the do notation emulates in Haskell.

You might be wondering then, why use monads at all? If we want to make side effects invisible, why not stick to an imperative language? The answer is that the monad gives us much better control over side effects. For instance, the log in the Writer monad is passed from function to function and is never exposed globally. There is no possibility of garbling the log or creating a data race. Also, monadic code is clearly demarcated and cordoned off from the rest of the program.

The do notation is just syntactic sugar for monadic composition. On the surface, it looks a lot like imperative code, but it translates directly to a sequence of binds and lambda expressions.

For instance, take the example we used previously to illustrate the composition of Kleisli arrows in the Writer monad. Using our current definitions, it could be rewritten as:

process :: String -> Writer String [String]
process = upCase >=> toWords

This function turns all characters in the input string to upper case and splits it into words, all the while producing a log of its actions.

In the do notation it would look like this:

process s = do
    upStr <- upCase s
    toWords upStr

Here, upStr is just a String, even though upCase produces a Writer:

upCase :: String -> Writer String String
upCase s = Writer (map toUpper s, "upCase ")

This is because the do block is desugared by the compiler to:

process s = 
   upCase s >>= \ upStr ->
       toWords upStr

The monadic result of upCase is bound to a lambda that takes a String. It’s the name of this string that shows up in the do block. When reading the line:

upStr <- upCase s

we say that upStr gets the result of upCase s.

The pseudo-imperative style is even more pronounced when we inline toWords. We replace it with the call to tell, which logs the string "toWords ", followed by the call to return with the result of splitting the string upStr using words. Notice that words is a regular function working on strings.

process s = do
    upStr <- upStr s
    tell "toWords "
    return (words upStr)

Here, each line in the do block introduces a new nested bind in the desugared code:

process s = 
    upCase s >>= \upStr ->
      tell "toWords " >>= \() ->
        return (words upStr)

Notice that tell produces a unit value, so it doesn’t have to be passed to the following lambda. Ignoring the contents of a monadic result (but not its effect — here, the contribution to the log) is quite common, so there is a special operator to replace bind in that case:

(>>) :: m a -> m b -> m b
m >> k = m >>= (\_ -> k)

The actual desugaring of our code looks like this:

process s = 
    upCase s >>= \upStr ->
      tell "toWords " >>
        return (words upStr)

In general, do blocks consist of lines (or sub-blocks) that either use the left arrow to introduce new names that are then available in the rest of the code, or are executed purely for side-effects. Bind operators are implicit between the lines of code. Incidentally, it is possible, in Haskell, to replace the formatting in the do blocks with braces and semicolons. This provides the justification for describing the monad as a way of overloading the semicolon.

Notice that the nesting of lambdas and bind operators when desugaring the do notation has the effect of influencing the execution of the rest of the do block based on the result of each line. This property can be used to introduce complex control structures, for instance to simulate exceptions.

Interestingly, the equivalent of the do notation has found its application in imperative languages, C++ in particular. I’m talking about resumable functions or coroutines. It’s not a secret that C++ futures form a monad. It’s an example of the continuation monad, which we’ll discuss shortly. The problem with continuations is that they are very hard to compose. In Haskell, we use the do notation to turn the spaghetti of “my handler will call your handler” into something that looks very much like sequential code. Resumable functions make the same transformation possible in C++. And the same mechanism can be applied to turn the spaghetti of nested loops into list comprehensions or “generators,” which are essentially the do notation for the list monad. Without the unifying abstraction of the monad, each of these problems is typically addressed by providing custom extensions to the language. In Haskell, this is all dealt with through libraries.

Next: Monads and Effects.

In the previous blog post we talked about relations. I gave an example of a thin category as a kind of relation that’s compatible with categorical structure. In a thin category, the hom-set is either an empty set or a singleton set. It so happens that these two sets form a cub-category of Set. It’s a very interesting category. It consists of the two objects — let’s give them new names o and i. Besides the mandatory identity morphisms, we also have a single morphism going from o to i, corresponding to the function we call absurd in Haskell:

absurd :: Void -> a
absurd _ = a

This tiny category is sometimes called the interval category. I’ll call it o->i.

Impoverished 4

The object o is initial, and the object i is terminal — just as the empty set and the singleton set were in Set. Moreover, the cartesian product from Set can be used to define a tensor product in o->i. We’ll use this tensor product to build a monoidal category.

Monoidal Categories

A tensor product is a bifunctor ⊗ with some additional properties. Here, in the interval category, we’ll define it through the following multiplication table:

o ⊗ o = o
o ⊗ i = o
i ⊗ o = o
i ⊗ i = i

Its action on pairs of morphisms (what we call bimap in Haskell) is also easy to define. For instance, what’s the action of on the pair <absurd, idi>? This pair takes the pair <o, i> to <i, i>. Under the bifunctor , the first pair produces o, and the second i. There is only one morphism from o to i, so we have:

absurd ⊗ idi = absurd

If we designate the (terminal) object i as the unit of the tensor product, we get a (symmetric) monoidal category. A monoidal category is a category with a tensor product that’s associative and unital (usually, up to isomorphism — but here, strictly).

Now imagine that we replace hom-sets in our original thin category with objects from the monoidal category o->i (we’ll call them hom-objects). After all, we were only using two sets from Set. We can replace the empty hom-set with the object o, and the singleton hom-set with the object i. We get what’s called an enriched category (although, in this case, it’s more of an impoverished category).

Impoverished 9

An example of a thin category (a total order with objects 1, 2, and 3) with hom-sets replaced by hom-objects from the interval category. Think of i as corresponding to less-than-or-equal, and o as greater.

Enriched Categories

An enriched category has hom-objects instead of hom-sets. These are objects from some monoidal category V called the base category. The base category has to be monoidal because we want to define something that would replace the usual composition of morphisms. Morphisms are elements of hom-sets. However, hom-objects, in general, have no elements. We don’t know what an element of o or i is.

So to fully define an enriched category we have to come up with a sensible substitute for composition. To do that, we need to rethink composition — first in terms of hom-sets, then in terms of hom-objects.

We can think of composition as a function from a cartesian product of two hom-sets to a third hom-set:

composea b c :: C(b, c) × C(a, b) -> C(a, c)

Generalizing it, we can replace hom-sets with hom-objects (here, either o or i), the cartesian product with the tensor product, and a function with a morphism (notice: it’s a morphism in our monoidal category o->i). These composition-defining morphisms form a “composition table” for hom-objects.

As an example, take the composition of two is. Their product i ⊗ i is i again, and there is only one morphism out of i, the identity morphism. In terms of original hom-sets it would mean that the composition of two morphisms always exists. In general, we have to impose this condition when we’re defining a category, enriched or not — here it just happens automatically.

For instance (see illustration), compose0 1 2=idi:

compose0 1 2 (C(1, 2) ⊗ C(0, 1))
= compose0 1 2 (i ⊗ i)
= compose0 1 2 i
= i
= C(0, 2)

In every category we must also have identity morphisms. These are special elements in the hom-sets of the form C(a, a). We have to find a way to define their equivalent in the enriched setting. We’ll use the standard trick of defining generalized elements. It’s based on the observation that selecting an element from a set s is the same as selecting a morphism that goes from the singleton set (the terminal object in Set) to s. In a monoidal category, we replace the terminal object with the monoidal unit.

So, instead of picking an identity morphism in C(a, a), we use a morphism from the monoidal unit i:

ja :: i -> C(a, a)

Again, in the case of a thin category, there is only one morphism leaving i, and that’s the identity morphism. That’s why we are automatically guaranteed that, in a thin category, all hom-objects of the form C(a, a) are equal to i.

Composition in a category must also satisfy associativity and identity conditions. Associativity in the enriched setting translates straightforwardly to a commuting diagram, but identity is a little trickier. We have to use ja to “select” the identity from the hom-object C(a, a) while composing it with some other hom-object C(b, a). We start with the product:

i ⊗ C(b, a)

Because i is the monoidal unit, this is equal to C(b, a). On the other hand, we can tensor together two morphisms in o->i — remember, a tensor product is a bifunctor, so it also acts on morphisms. Here we’ll tensor ja and the identity at C(b, a):

ja ⊗ idC(b, a)

We act with this product on the product object i ⊗ C(b, a) to get C(a, a) ⊗ C(b, a). Then we use composition to get:

C(a, a) ⊗ C(b, a) -> C(b, a)

These two ways of getting to C(b, a) must coincide, leading to the identity condition for enriched categories.

Impoverished 5

Now that we’ve seen how the enrichment works for thin categories, we can apply the same mechanism to define categories enriched over any monoidal category V.

The important part is that V defines a (bifunctor) tensor product ⊗ and a unit object i. Associativity and unitality may be either strict or up to isomorphism (notice that a regular cartesian product is associative only up to isomorphism — (a, (b, c)) is not equal to ((a, b), c)).

Instead of sets of morphisms, an enriched category has hom-objects that are objects in V. We use the same notation as for hom-sets: C(a, b) is the hom-object that connects object a to object b. Composition is replaced by morphisms in V:

composea b c :: C(b, c) ⊗ C(a, b) -> C(a, c)

Instead of identity morphisms, we have the morphisms in V:

ja :: i -> C(a, a)

Finally, associativity and unitality of composition are imposed in the form of a few commuting diagrams.

Impoverished Yoneda

The Yoneda Lemma talks about functors from an arbitrary category to Set. To generalize the Yoneda lemma to enriched categories we first have to generalize functors. Their action on objects is not a problem; it’s the action on morphisms that needs our attention.

Enriched Functors

Since in an enriched category we no longer have access to individual morphisms, we have to define the action of functors on hom-objects wholesale. This is only possible if the hom-objects in the target category come from the same base category V as the hom-objects in the source category. In other words, both categories must be enriched over the same monoidal category. We can then use regular morphisms in V to map hom-objects.

Between any two objects a and b in C we have the hom-object C(a, b). The two objects are mapped by the functor f to f a and f b, and there is a hom-object between them, D(f a, f b). The action of f on C(a, b) is defined as a morphism in V:

C(a, b) -> D(f a, f b)

Impoverished 6

Let’s see what this means in our impoverished thin category. First of all, a functor will always map related objects to related objects. That’s because there is no morphism from i to o. A bond between two objects cannot be broken by an impoverished functor.

If the relation is a partial order, for instance less-than-or-equal, then it follows that a functor between posets preserves the ordering — it’s monotone.

A functor must also preserve composition and identity. The former can be easily expressed as a commuting diagram. Identity preservation in the enriched setting involves the use of ja. Starting from i we can use ja to get to C(a, a), which the functor maps to D(f a, f a). Or we can use jf a to get there directly. We insist that both paths be the same.

Impoverished 7

In our impoverished category, this just works because ja is the identity morphism and all C(a, a)s and D(a, a)s are equal to i.

Back to Yoneda: You might remember that we start the Yoneda construction by fixing one object a in C, and then varying another object x to define the functor:

x -> C(a, x)

This functor maps C to Set, because xs are objects in C, and hom-sets are sets — objects of Set.

In the enriched environment, the same construction results in a mapping from C to V, because hom-objects are objects of the base category V.

But is this mapping a functor? This is far from obvious, considering that C is an enriched category, and we have just said that enriched functors can only go between categories that are enriched over the same base category. The target of our functor, the category V, is not enriched. It turns out that, as long as V is closed, we can turn it into an enriched category.

Self Enrichment

Let’s first see how we would enrich our tiny category o->i. First of all, let’s check if it’s closed. Closedness means that hom-sets can be objectified — for every hom-set there is an object called the exponential object that objectifies it. The exponential object in a (symmetric) monoidal category is defined through the adjunction:

V(a⊗b, c) ≅ V(b, ca)

This is the standard adjunction for defining exponentials, except that we are using the tensor product instead of the regular product. The hom-sets are sets of morphisms between objects in V (here, in o->i).

Let’s check, for instance, if there’s an object that corresponds to the hom-set V(o, i), which we would call io. We have:

V(o⊗b, i) ≅ V(b, io)

Whatever b we chose, when multiplied by o it will yield o, so the left hand side is V(o, i), a singleton set. Therefore V(b, io) must be a singleton set too, for any choice of b. In particular, if b is i, we see that the only choice for io is:

io = i

You can check that all exponentiation rules in o->i can be obtained from simple algebra by replacing o with zero and i with one.

Every closed symmetric monoidal category can be enriched in itself by replacing hom-sets with the corresponding exponentials. For instance, in our case, we end up replacing all empty hom-sets in the category o->i with o, and all singleton hom-sets with i. You can easily convince yourself that it works, and the result is the category o->i enriched in itself.

Impoverished 8

We can now take a category C that’s enriched over a closed symmetric monoidal category V, and show that the mapping:

x -> C(a, x)

is indeed an enriched functor. It maps objects of C to objects of V and hom-objects of C to hom-objects (exponentials) of V.

Impoverished 10

An example of a functor from a total order enriched over the interval category to the interval category. This particular functor is equal to the hom-functor C(a->x) for a equal to 3.

Let’s see what this functor looks like in a poset. Given some a, the hom-object C(a, x) is equal to i if a <= x. So an x is mapped to i if it’s greater-or-equal to a, otherwise it’s mapped to o. If you think of the objects mapped to o as colored black and the ones mapped to i as colored red, you’ll see the object a and the whole graph below it must be painted red.

Enriched Natural Transformations

Now that we know what enriched functors are, we have to define natural transformations between them. This is a bit tricky, since a regular natural transformation is defined as a family of morphisms. But again, instead of picking individual morphisms from hom-sets we can work with the closest equivalent: generalized elements — morphisms going from the unit object i to hom-objects. So an enriched natural transformation between two enriched functors f and g is defined as a family of morphisms in V:

αa :: i -> V(f a, g a)

Natural transformations are very limited in our impoverished category. Let’s see what morphisms from i are at our disposal. We have one morphism from i to i: the identity morphism ida. This makes sense — we think of i as having a single element. There is no morphism from i back to o; and that makes sense too — we think of o as having no elements. The only possible generalized components of an impoverished natural transformation between two functors f and g correspond to D(f a, g a) equal to i; which means that, for every a, f a must be less-than-or-equal to g a. A natural transformation can only push a functor uphill.

When the target category is o->i, as in the impoverished Yoneda lemma, a natural transformation may never connect red to black. So once the first functor switches to red, the other must follow.

Naturality Condition

There is, of course, a naturality condition that goes with this definition of a natural transformation. The essence of it is that it shouldn’t matter if we first apply a functor and then the natural transformation α, or the other way around. In the enriched context, there are two ways of getting from C(a, b) to D(f a, g b). One is to multiply C(a, b) by i on the right:

C(a, b) ⊗ i

apply the product of g ⊗ αa to get:

D(g a, g b) ⊗ D(f a, g a)

and then apply composition to get:

D(f a, g b)

The other way is to multiply C(a, b) by i on the left:

i ⊗ C(a, b)

apply αb ⊗ f to get:

D(f b, g b) ⊗ D(f a, f b)

and compose the two to get:

D(f a, g b)

The naturality condition requires that this diagram commute.

Impoverished 11

Enriched Yoneda

The enriched version of the Yoneda lemma talks about enriched natural transformations from the functor x -> C(a, x) to any enriched functor f that goes from C to V.

Consider for a moment a functor from a poset to our tiny category o->i (which, by the way, is also a poset). It will map some objects to o (black) and others to i (red). As we’ve seen, a functor must preserve the less-than-or-equal relation, so once we get into the red territory, there is no going back to black. And a natural transformation may only repaint black to red, not the other way around.

Now we would like to say that natural transformations from x -> C(a, x) to f are in one-to-one correspondence with the elements of f a, except that f a is not a set, so it doesn’t have elements. It’s an object in V. So instead of talking about elements of f a, we’ll talk about generalized elements — morphisms from the unit object i to f a. And that’s how the enriched Yoneda lemma is formulated — as a natural bijection between the set of natural transformations and a set of morphisms from the unit object to f a.

Nat(C(a, -), f) ≅ i -> f a

In our running example, there are only two possible values for f a.

  1. If the value is o then there is no morphism from i to it. The Yoneda lemma tells us that there is no natural transformation in that case. That makes sense, because the value of the functor x -> C(a, x) at x=a is i, and there is no morphism from i to o.
  2. If the value is i then there is exactly one morphism from i to it — the identity. The Yoneda lemma tells use that there is just one natural transformation in that case. It’s the natural transformation whose generalized component at any object x is i->i.

Strong Enriched Yoneda

There is something unsatisfactory in the fact that the enriched Yoneda lemma ends up using a mapping between sets. First we try to get away from sets as far as possible, then we go back to sets of morphisms. It feels like cheating. Not to worry! There is a stronger version of the Yoneda lemma that deals with this problem. What we need is to replace the set of natural transformations with an object in V that would represent them — just like we replaced the set of morphisms with the exponential object. Such an object is defined as an end:

x V(f x, g x)

The strong version of the Yoneda lemma establishes the natural isomorphism:

x V(C(a, x), f x) ≅ f a

Enriched Profunctors

We’ve seen that a profunctor is a functor from a product category Cop × D to Set. The enriched version of a profunctor requires the notion of a product of enriched categories. We would like the product of enriched categories to also be an enriched category. In fact, we would like it to be enriched over the same base category V as the component categories.

We’ll define objects in such a category as pairs of objects from the component categories, but the hom-objects will be defined as tensor products of the component hom-objects. In the enriched product category, the hom-object between two pairs, <c, d> and <c', d'> is:

(Cop ⊗ D)(<c, d>, <c', d'>) = C(c, c') ⊗ D(d, d')

You can convince yourself that composition of such hom-objects requires the tensor product to be symmetric (at least up to isomorphism). That’s because you have to be able to rearrange the hom-objects in a tensor product of tensor products.

An enriched profunctor is defined as an enriched functor from the tensor product of two categories to the (self-enriched) base category:

Cop ⊗ D -> V

Just like regular profunctors, enriched profunctors can be composed using the coend formula. The only difference is that the cartesian product is replaced by the tensor product in V. They form a bicategory called V-Prof.

Enriched profunctors are the basis of the definition of Tambara modules, which are relevant in the application to Haskell lenses.


One of the reasons for using category theory is to get away from set theory. In general, objects in a category don’t have to form sets. The morphisms, however, are elements of sets — the hom-sets. Enriched categories go a step further and replace even those sets with categorical objects. However, it’s not categories all the way down — the base category that’s used for enrichment is still a regular old category with hom-sets.


I’m grateful to Gershom Bazerman for useful comments and to André van Meulebrouck for checking the grammar and spelling.

This is part 18 of Categories for Programmers. Previously: It’s All About Morphisms. See the Table of Contents.

In mathematics we have various ways of saying that one thing is like another. The strictest is equality. Two things are equal if there is no way to distinguish one from another. One can be substituted for the other in every imaginable context. For instance, did you notice that we used equality of morphisms every time we talked about commuting diagrams? That’s because morphisms form a set (hom-set) and set elements can be compared for equality.

But equality is often too strong. There are many examples of things being the same for all intents and purposes, without actually being equal. For instance, the pair type (Bool, Char) is not strictly equal to (Char, Bool), but we understand that they contain the same information. This concept is best captured by an isomorphism between two types — a morphism that’s invertible. Since it’s a morphism, it preserves the structure; and being “iso” means that it’s part of a round trip that lands you in the same spot, no matter on which side you start. In the case of pairs, this isomorphism is called swap:

swap       :: (a,b) -> (b,a)
swap (a,b) = (b,a)

swap happens to be its own inverse.

Adjunction and Unit/Counit Pair

When we talk about categories being isomorphic, we express this in terms of mappings between categories, a.k.a. functors. We would like to be able to say that two categories C and D are isomorphic if there exists a functor R (“right”) from C to D, which is invertible. In other words, there exists another functor L (“left”) from D back to C which, when composed with R, is equal to the identity functor I. There are two possible compositions, R ∘ L and L ∘ R; and two possible identity functors: one in C and another in D.

Adj - 1

But here’s the tricky part: What does it mean for two functors to be equal? What do we mean by this equality:

R ∘ L = ID

or this one:

L ∘ R = IC

It would be reasonable to define functor equality in terms of equality of objects. Two functors, when acting on equal objects, should produce equal objects. But we don’t, in general, have the notion of object equality in an arbitrary category. It’s just not part of the definition. (Going deeper into this rabbit hole of “what equality really is,” we would end up in Homotopy Type Theory.)

You might argue that functors are morphisms in the category of categories, so they should be equality-comparable. And indeed, as long as we are talking about small categories, where objects form a set, we can indeed use the equality of elements of a set to equality-compare objects.

But, remember, Cat is really a 2-category. Hom-sets in a 2-category have additional structure — there are 2-morphisms acting between 1-morphisms. In Cat, 1-morphisms are functors, and 2-morphisms are natural transformations. So it’s more natural (can’t avoid this pun!) to consider natural isomorphisms as substitutes for equality when talking about functors.

So, instead of isomorphism of categories, it makes sense to consider a more general notion of equivalence. Two categories C and D are equivalent if we can find two functors going back and forth between them, whose composition (either way) is naturally isomorphic to the identity functor. In other words, there is a two-way natural transformation between the composition R ∘ L and the identity functor ID, and another between L ∘ R and the identity functor IC.

Adjunction is even weaker than equivalence, because it doesn’t require that the composition of the two functors be isomorphic to the identity functor. Instead it stipulates the existence of a one way natural transformation from ID to R∘L, and another from L∘R to IC. Here are the signatures of these two natural transformations:

η :: ID -> R ∘ L
ε :: L ∘ R -> IC

η is called the unit, and ε the counit of the adjunction.

Notice the asymmetry between these two definitions. In general, we don’t have the two remaining mappings:

R ∘ L -> ID -- not necessarily
IC -> L ∘ R -- not necessarily

Because of this asymmetry, the functor L is called the left adjoint to the functor R, while the functor R is the right adjoint to L. (Of course, left and right make sense only if you draw your diagrams one particular way.)

The compact notation for the adjunction is:

L ⊣ R

To better understand the adjunction, let’s analyze the unit and the counit in more detail.


Let’s start with the unit. It’s a natural transformation, so it’s a family of morphisms. Given an object d in D, the component of η is a morphism between I d, which is equal to d, and (R ∘ L) d; which, in the picture, is called d':

ηd :: d -> (R ∘ L) d

Notice that the composition R∘L is an endofunctor in D.

This equation tells us that, to get from object d to d' we can either take the shortcut using the morphism ηd in D, or do the round trip using two functors, with a layover in C. Or, stretching the analogy, we can take the airport shuttle from Departures to Arrivals at the airport D and skip the flights altogether. Notice that the morphism is a one way trip. There isn’t necessarily a shuttle going from d' back to d


By the same token, the component of of the counit ε can be described as:

εc' :: (L ∘ R) c -> c

where c' is (L ∘ R) c. It tells us that, after we’ve gone on the round trip with a layover in D, we can take the airport shuttle to get back to our starting point.

Notice again the asymmetry between unit and counit. Unit saves you a round trip, while counit completes the round trip. (Useful mnemonic: counit completes.)

Another way of looking at unit and counit is that unit lets us introduce the composition R ∘ L anywhere we could insert an identity functor on D; and counit lets us eliminate the composition L ∘ R, replacing it with the identity on C. That leads to some “obvious” consistency conditions, which make sure that introduction followed by elimination doesn’t change anything:

R = ID ∘ R -> R ∘ L ∘ R -> R ∘ IC = R
L = L ∘ ID -> L ∘ R ∘ L -> IC ∘ L  = L

These are called triangular identities because they make the following diagrams commute (these are diagrams in the functor category):



We often see unit and counit in Haskell under different names. Unit is known as return (or pure, in the definition of Applicative):

return :: d -> m d

and counint as extract:

extract :: w c -> c

Here, m is the (endo-) functor corresponding to R∘L, and w is the (endo-) functor corresponding to L∘R. As we’ll see later, they are part of the definition of a monad and a comonad, respectively.

If you think of an endofunctor as a container, the unit (or return) is a polymorphic function that creates a default box around a value of arbitrary type. The counit (or extract) does the reverse: it retrieves or produces a single value from a container.

We’ll see later that every pair of adjoint functors defines a monad and a comonad. Conversely, every monad or comonad may be factorized into a pair of adjoint functors — this factorization is not unique, though.

In Haskell, we use monads a lot, but only rarely factorize them into pairs of adjoint functors, primarily because those functors would normally take us out of Hask.

We can however define adjunctions of endofunctors in Haskell. Here’s part of the definition taken from Data.Functor.Adjunction:

class (Functor f, Representable u) =>
      Adjunction f u | f -> u, u -> f where
  unit         :: a -> u (f a)
  counit       :: f (u a) -> a

This definition requires some explanation. First of all, it describes a multi-parameter type class — the two parameters being f and u. It establishes a relation called Adjunction between these two type constructors.

Additional conditions, after the vertical bar, specify functional dependencies. For instance, f -> u means that u is determined by f (the relation between f and u is a function, here on type constructors). Conversely, u -> f means that, if we know u, f is uniquely determined.

I’ll explain in a moment why, in Haskell, we can impose the condition that the right adjoint u be a representable functor.

Adjunction and Hom-Sets

There is an equivalent definition of the adjunction in terms of natural isomorphisms of hom-sets. This definition ties nicely with universal constructions we’ve been studying so far. Every time you hear the statement that there is some unique morphism, which factorizes some construction, you should think of it as a mapping of some set to a hom-set. That’s the meaning of “picking a unique morphism.”

Furthermore, factorization can be often described in terms of natural transformations. Factorization involves commuting diagrams — some morphism being equal to a composition of two morphisms (factors). A natural transformation maps morphisms to commuting diagrams. So, in a universal construction, we go from a morphism to a commuting diagram, and then to a unique morphism. We end up with a mapping from morphism to morphism, or from one hom-set to another (usually in different categories). If this mapping is invertible, and if it can be naturally extended across all hom-sets, we have an adjunction.

The main difference between universal constructions and adjunctions it that the latter are defined globally — for all hom-sets. For instance, using a universal construction you can define a product of two select objects, even if it doesn’t exist for any other pair of objects in that category. As we’ll see soon, if the product of any pair of objects exists in a category, it can be also defined through an adjunction.


Here’s the alternative definition of the adjunction using hom-sets. As before, we have two functors L :: D->C and R :: C->D. We pick two arbitrary objects: the source object d in D, and the target object c in C. We can map the source object d to C using L. Now we have two objects in C, L d and c. They define a hom-set:

C(L d, c)

Similarly, we can map the target object c using R. Now we have two objects in D, d and R c. They, too, define a hom set:

D(d, R c)

We say that L is left adjoint to R iff there is an isomorphism of hom sets:

C(L d, c) ≅ D(d, R c)

that is natural both in d and c.

Naturality means that the source d can be varied smoothly across D; and the target c, across C. More precisely, we have a natural transformation φ between the following two (covariant) functors from C to Set. Here’s the action of these functors on objects:

c -> C(L d, c)
c -> D(d, R c)

The other natural transformation, ψ, acts between the following (contravariant) functors:

d -> C(L d, c)
d -> D(d, R c)

Both natural transformations must be invertible.

It’s easy to show that the two definitions of the adjunction are equivalent. For instance, let’s derive the unit transformation starting from the isomorphism of hom-sets:

C(L d, c) ≅ D(d, R c)

Since this isomorphism works for any object c, it must also work for c = L d:

C(L d, L d) ≅ D(d, (R ∘ L) d)

We know that the left hand side must contain at least one morphism, the identity. The natural transformation will map this morphism to an element of D(d, (R ∘ L) d) or, inserting the identity functor I, a morphism in:

D(I d, (R ∘ L) d)

We get a family of morphisms parameterized by d. They form a natural transformation between the functor I and the functor R ∘ L (the naturality condition is easy to verify). This is exactly our unit, η.

Conversely, starting from the existence of the unit and co-unit, we can define the transformations between hom-sets. For instance, let’s pick an arbitrary morphism f in the hom-set C(L d, c). We want to define a φ that, acting on f, produces a morphism in D(d, R c).

There isn’t really much choice. One thing we can try is to lift f using R. That will produce a morphism R f from R (L d) to R c — a morphism that’s an element of D((R ∘ L) d, R c).

What we need for a component of φ, is a morphism from d to R c. That’s not a problem, since we can use a component of ηd to get from d to (R ∘ L) d. We get:

φf = R f ∘ ηd

The other direction is analogous, and so is the derivation of ψ.

Going back to the Haskell definition of Adjunction, the natural transformations φ and ψ are replaced by polymorphic (in a and b) functions leftAdjunct and rightAdjunct, respectively. The functors L and R are called f and u:

class (Functor f, Representable u) =>
      Adjunction f u | f -> u, u -> f where
  leftAdjunct  :: (f a -> b) -> (a -> u b)
  rightAdjunct :: (a -> u b) -> (f a -> b)

The equivalence between the unit/counit formulation and the leftAdjunct/rightAdjunct formulation is witnessed by these mappings:

  unit           = leftAdjunct id
  counit         = rightAdjunct id
  leftAdjunct f  = fmap f . unit
  rightAdjunct f = counit . fmap f

It’s very instructive to follow the translation from the categorical description of the adjunction to Haskell code. I highly encourage this as an exercise.

We are now ready to explain why, in Haskell, the right adjoint is automatically a representable functor. The reason for this is that, to the first approximation, we can treat the category of Haskell types as the category of sets.

When the right category D is Set, the right adjoint R is a functor from C to Set. Such a functor is representable if we can find an object rep in C such that the hom-functor C(rep, _) is naturally isomorphic to R. It turns out that, if R is the right adjoint of some functor L from Set to C, such an object always exists — it’s the image of the singleton set () under L:

rep = L ()

Indeed, the adjunction tells us that the following two hom-sets are naturally isomorphic:

C(L (), c) ≅ Set((), R c)

For a given c, the right hand side is the set of functions from the singleton set () to R c. We’ve seen earlier that each such function picks one element from the set R c. The set of such functions is isomorphic to the set R c. So we have:

C(L (), -) ≅ R

which shows that R is indeed representable.

Product from Adjunction

We have previously introduced several concepts using universal constructions. Many of those concepts, when defined globally, are easier to express using adjunctions. The simplest non-trivial example is that of the product. The gist of the universal construction of the product is the ability to factorize any product-like candidate through the universal product.

More precisely, the product of two objects a and b is the object (a × b) (or (a, b) in the Haskell notation) equipped with two morphisms fst and snd such that, for any other candidate c equipped with two morphisms p::c->a and q::c->b, there exists a unique morphism m::c->(a, b) that factorizes p and q through fst and snd.

As we’ve seen earlier, in Haskell, we can implement a factorizer that generates this morphism from the two projections:

factorizer :: (c -> a) -> (c -> b) -> (c -> (a, b))
factorizer p q = \x -> (p x, q x)

It’s easy to verify that the factorization conditions hold:

fst . factorizer p q = p
snd . factorizer p q = q

We have a mapping that takes a pair of morphisms p and q and produces another morphism m = factorizer p q.

How can we translate this into a mapping between two hom-sets that we need to define an adjunction? The trick is to go outside of Hask and treat the pair of morphisms as a single morphism in the product category.

Let me remind you what a product category is. Take two arbitrary categories C and D. The objects in the product category C×D are pairs of objects, one from C and one from D. The morphisms are pairs of morphisms, one from C and one from D.

To define a product in some category C, we should start with the product category C×C. Pairs of morphism from C are single morphisms in the product category C×C.


It might be a little confusing at first that we are using a product category to define a product. These are, however, very different products. We don’t need a universal construction to define a product category. All we need is the notion of a pair of objects and a pair of morphisms.

However, a pair of objects from C is not an object in C. It’s an object in a different category, C×C. We can write the pair formally as <a, b>, where a and b are objects of C. The universal construction, on the other hand, is necessary in order to define the object a×b (or (a, b) in Haskell), which is an object in the same category C. This object is supposed to represent the pair <a, b> in a way specified by the universal construction. It doesn’t always exist and, even if it exists for some, might not exist for other pairs of objects in C.

Let’s now look at the factorizer as a mapping of hom-sets. The first hom-set is in the product category C×C, and the second is in C. A general morphism in C×C would be a pair of morphisms <f, g>:

f :: c' -> a
g :: c'' -> b

with c'' potentially different from c'. But to define a product, we are interested in a special morphism in C×C, the pair p and q that share the same source object c. That’s okay: In the definition of an adjuncion, the source of the left hom-set is not an arbitrary object — it’s the result of the left functor L acting on some object from the right category. The functor that fits the bill is easy to guess — it’s the diagonal functor from C to C×C, whose action on objects is:

Δ c = <c, c>

The left-hand side hom-set in our adjunction should thus be:

(C×C)(Δ c, <a, b>)

It’s a hom-set in the product category. Its elements are pairs of morphisms that we recognize as the arguments to our factorizer:

(c -> a) -> (c -> b) ...

The right-hand side hom-set lives in C, and it goes between the source object c and the result of some functor R acting on the target object in C×C. That’s the functor that maps the pair <a, b> to our product object, a×b. We recognize this element of the hom-set as the result of the factorizer:

... -> (c -> (a, b))


We still don’t have a full adjunction. For that we first need our factorizer to be invertible — we are building an isomorphism between hom-sets. The inverse of the factorizer should start from a morphism m — a morphism from some object c to the product object a×b. In other words, m should be an element of:

C(c, a×b)

The inverse factorizer should map m to a morphism <p, q> in C×C that goes from <c, c> to <a, b>; in other words, a morphism that’s an element of:

(C×C)(Δ c, <a, b>)

If that mapping exists, we conclude that there exists the right adjoint to the diagonal functor. That functor defines a product.

In Haskell, we can always construct the inverse of the factorizer by composing m with, respectively, fst and snd.

p = fst ∘ m
q = snd ∘ m

To complete the proof of the equivalence of the two ways of defining a product we also need to show that the mapping between hom-sets is natural in a, b, and c. I will leave this as an exercise for the dedicated reader.

To summarize what we have done: A categorical product may be defined globally as the right adjoint of the diagonal functor:

(C × C)(Δ c, <a, b>) ≅ C(c, a×b)

Here, a×b is the result of the action of our right adjoint functor Product on the pair <a, b>. Notice that any functor from C×C is a bifunctor, so Product is a bifunctor. In Haskell, the Product bifunctor is written simply as (,). You can apply it to two types and get their product type, for instance:

(,) Int Bool ~ (Int, Bool)

Exponential from Adjunction

The exponential ba, or the function object a⇒b, can be defined using a universal construction. This construction, if it exists for all pairs of objects, can be seen as an adjunction. Again, the trick is to concentrate on the statement:

For any other object z with a morphism

g :: z × a -> b

there is a unique morphism

h :: z -> (a⇒b)

This statement establishes a mapping between hom-sets.

In this case, we are dealing with objects in the same category, so the two adjoint functors are endofunctors. The left (endo-)functor L, when acting on object z, produces z × a. It’s a functor that corresponds to taking a product with some fixed a.

The right (endo-)functor R, when acting on b produces the function object a⇒b (or ba). Again, a is fixed. The adjunction between these two functors is often written as:

- × a ⊣ (-)a

The mapping of hom-sets that underlies this adjunction is best seen by redrawing the diagram that we used in the universal construction.


Notice that the eval morphism is nothing else but the counit of this adjunction:

(a⇒b) × a -> b


(a⇒b) × a = (L ∘ R) b

I have previously mentioned that a universal construction defines a unique object, up to isomorphism. That’s why we have “the” product and “the” exponential. This property translates to adjunctions as well: if a functor has an adjoint, this adjoint is unique up to isomorphism.


  1. Derive the naturality square for ψ, the transformation between the two (contravariant) functors:
    a -> C(L a, b)
    a -> D(a, R b)
  2. Derive the counit ε starting from the hom-sets isomorphism in the second definition of the adjunction.
  3. Complete the proof of equivalence of the two definitions of the adjunction.
  4. Show that the coproduct can be defined by an adjunction. Start with the definition of the factorizer for a coproduct.
  5. Show that the coproduct is the left adjoint of the diagonal functor.
  6. Define the adjunction between a product and a function object in Haskell.

Next: Free/Forgetful Adjunctions.


I’d like to thank Edward Kmett and Gershom Bazerman for checking my math and logic, and André van Meulebrouck, who has been volunteering his editing help throughout this series of posts.

This is part 17 of Categories for Programmers. Previously: Yoneda Embedding. See the Table of Contents.

If I haven’t convinced you yet that category theory is all about morphisms then I haven’t done my job properly. Since the next topic is adjunctions, which are defined in terms of isomorphisms of hom-sets, it makes sense to review our intuitions about the building blocks of hom-sets. Also, you’ll see that adjunctions provide a more general language to describe a lot of constructions we’ve studied before, so it might help to review them too.


To begin with, you should really think of functors as mappings of morphisms — the view that’s emphasized in the Haskell definition of the Functor typeclass, which revolves around fmap. Of course, functors also map objects — the endpoints of morphisms — otherwise we wouldn’t be able to talk about preserving composition. Objects tell us which pairs of morphisms are composable. The target of one morphism must be equal to the source of the other — if they are to be composed. So if we want the composition of morphisms to be mapped to the composition of lifted morphisms, the mapping of their endpoints is pretty much determined.

Commuting Diagrams

A lot of properties of morphisms are expressed in terms of commuting diagrams. If a particular morphism can be described as a composition of other morphisms in more than one way, then we have a commuting diagram.

In particular, commuting diagrams form the basis of almost all universal constructions (with the notable exceptions of the initial and terminal objects). We’ve seen this in the definitions of products, coproducts, various other (co-)limits, exponential objects, free monoids, etc.

The product is a simple example of a universal construction. We pick two objects a and b and see if there exists an object c, together with a pair of morphisms p and q, that has the universal property of being their product.


A product is a special case of a limit. A limit is defined in terms of cones. A general cone is built from commuting diagrams. Commutativity of those diagrams may be replaced with a suitable naturality condition for the mapping of functors. This way commutativity is reduced to the role of the assembly language for the higher level language of natural transformations.

Natural Transformations

In general, natural transformations are very convenient whenever we need a mapping from morphisms to commuting squares. Two opposing sides of a naturality square are the mappings of some morphism f under two functors F and G. The other sides are the components of the natural transformation (which are also morphisms).


Naturality means that when you move to the “neighboring” component (by neighboring I mean connected by a morphism), you’re not going against the structure of either the category or the functors. It doesn’t matter whether you first use a component of the natural transformation to bridge the gap between objects, and then jump to its neighbor using the functor; or the other way around. The two directions are orthogonal. A natural transformation moves you left and right, and the functors move you up and down or back and forth — so to speak. You can visualize the image of a functor as a sheet in the target category. A natural transformation maps one such sheet corresponding to F, to another, corresponding to G.


We’ve seen examples of this orthogonality in Haskell. There the action of a functor modifies the content of a container without changing its shape, while a natural transformation repackages the untouched contents into a different container. The order of these operations doesn’t matter.

We’ve seen the cones in the definition of a limit replaced by natural transformations. Naturality ensures that the sides of every cone commute. Still, a limit is defined in terms of mappings between cones. These mappings must also satisfy commutativity conditions. (For instance, the triangles in the definition of the product must commute.)

These conditions, too, may be replaced by naturality. You may recall that the universal cone, or the limit, is defined as a natural transformation between the (contravariant) hom-functor:

F :: c -> C(c, Lim D)

and the (also contravariant) functor that maps objects in C to cones, which themselves are natural transformations:

G :: c -> Nat(Δc, D)

Here, Δc is the constant functor, and D is the functor that defines the diagram in C. Both functors F and G have well defined actions on morphisms in C. It so happens that this particular natural transformation between F and G is an isomorphism.

Natural Isomorphisms

A natural isomorphism — which is a natural transformation whose every component is reversible — is category theory’s way of saying that “two things are the same.” A component of such a transformation must be an isomorphism between objects — a morphism that has the inverse. If you visualize functor images as sheets, a natural isomorphism is a one-to-one invertible mapping between those sheets.


But what are morphisms? They do have more structure than objects: unlike objects, morphisms have two ends. But if you fix the source and the target objects, the morphisms between the two form a boring set (at least for locally small categories). We can give elements of this set names like f or g, to distinguish one from another — but what is it, really, that makes them different?

The essential difference between morphisms in a given hom-set lies in the way they compose with other morphisms (from abutting hom-sets). If there is a morphism h whose composition (either pre- or post-) with f is different than that with g, for instance:

h ∘ f ≠ h ∘ g

then we can directly “observe” the difference between f and g. But even if the difference is not directly observable, we might use functors to zoom in on the hom-set. A functor F may map the two morphisms to distinct morphisms:

F f ≠ F g

in a richer category, where the abutting hom-sets provide more resolution, e.g.,

h' ∘ F f ≠ h' ∘ F g

where h' is not in the image of F.

Hom-Set Isomorphisms

A lot of categorical constructions rely on isomorphisms between hom-sets. But since hom-sets are just sets, a plain isomorphism between them doesn’t tell you much. For finite sets, an isomorphism just says that they have the same number of elements. If the sets are infinite, their cardinality must be the same. But any meaningful isomorphism of hom-sets must take into account composition. And composition involves more than one hom-set. We need to define isomorphisms that span whole collections of hom-sets, and we need to impose some compatibility conditions that interoperate with composition. And a natural isomorphism fits the bill exactly.

But what’s a natural isomorphism of hom-sets? Naturality is a property of mappings between functors, not sets. So we are really talking about a natural isomorphism between hom-set-valued functors. These functors are more than just set-valued functors. Their action on morphisms is induced by the appropriate hom-functors. Morphisms are canonically mapped by hom-functors using either pre- or post-composition (depending on the covariance of the functor).

The Yoneda embedding is one example of such an isomorphism. It maps hom-sets in C to hom-sets in the functor category; and it’s natural. One functor in the Yoneda embedding is the hom-functor in C and the other maps objects to sets of natural transformations between hom-sets.

The definition of a limit is also a natural isomorphism between hom-sets (the second one, again, in the functor category):

C(c, Lim D) ≃ Nat(Δc, D)

It turns out that our construction of an exponential object, or that of a free monoid, can also be rewritten as a natural isomorphism between hom-sets.

This is no coincidence — we’ll see next that these are just different examples of adjunctions, which are defined as natural isomorphisms of hom-sets.

Asymmetry of Hom-Sets

There is one more observation that will help us understand adjunctions. Hom-sets are, in general, not symmetric. A hom-set C(a, b) is often very different from the hom-set C(b, a). The ultimate demonstration of this asymmetry is a partial order viewed as a category. In a partial order, a morphism from a to b exists if and only if a is less than or equal to b. If a and b are different, then there can be no morphism going the other way, from b to a. So if the hom-set C(a, b) is non-empty, which in this case means it’s a singleton set, then C(b, a) must be empty, unless a = b. The arrows in this category have a definite flow in one direction.

A preorder, which is based on a relation that’s not necessarily antisymmetric, is also “mostly” directional, except for occasional cycles. It’s convenient to think of an arbitrary category as a generalization of a preoder.

A preorder is a thin category — all hom-sets are either singletons or empty. We can visualize a general category as a “thick” preorder.


  1. Consider some degenerate cases of a naturality condition and draw the appropriate diagrams. For instance, what happens if either functor F or G map both objects a and b (the ends of f :: a -> b) to the same object, e.g., F a = F b or G a = G b? (Notice that you get a cone or a co-cone this way.) Then consider cases where either F a = G a or F b = G b. Finally, what if you start with a morphism that loops on itself — f :: a -> a?

Next: Adjunctions.


I’d like to thank Gershom Bazerman for checking my math and logic, and André van Meulebrouck, who has been volunteering his editing help throughout this series of posts.

Next Page »