cat/doc/implementation.tex

\section{Implementation}
This implementation formalizes the following concepts:
%
\begin{itemize}
\item Core categorical concepts
\subitem Categories
\subitem Functors
\subitem Products
\subitem Exponentials
\subitem Cartesian closed categories
\subitem Natural transformations
\subitem Yoneda embedding
\subitem Monads
\subsubitem Monoidal monads
\subsubitem Kleisli monads
\subsubitem Voevodsky's construction
\item Category of \ldots
\subitem Homotopy sets
\subitem Categories
\subitem Relations
\subitem Functors
\subitem Free category
\end{itemize}
%
Since it is useful to distinguish between types with more or less (homotopical)
structure I have followed the following design-principle: I have split concepts
up into things that represent ``data'' and ``laws'' about this data. The idea is
that we can provide a proof that the laws are mere propositions. As an example a
category is defined to have two members: `raw` which is a collection of the data
and `isCategory` which asserts some laws about that data.

This allows me to reason about things in a more mathematical way, where one can
reason about two categories by simply focusing on the data. This is acheived by
creating a function embodying the ``equality principle'' for a given type.

\subsubsection{Categories}
The data for a category consist of objects, morphisms (or arrows as I will refer
to them henceforth), the identity arrow and composition of arrows.

Another record encapsulates some laws about this data: associativity of
composition, identity law for the identity morphism. These are standard
requirements for being a category as can be found in standard mathematical
expositions on the topic. We, however, impose one further requirement on what it
means to be a category, namely that the type of arrows form a set. We could
relax this requirement, this would give us the notion of higher categorier
(\cite[p. 307]{hott-2013}). For the purpose of this project, however, this
report will restrict itself to 1-categories.

Raw categories satisfying these properties are called a pre-categories.

As a further requirement to be a proper category we require it to be univalent.
This requirement is quite similiar to univalence for types, but we let
isomorphism of objects play the role of equivalence of types. The univalence
criterion is:
%
$$
\isEquiv\ (A \cong B)\ (A \equiv B)\ \idToIso_{A\ B}
$$
%
Note that this is a stronger requirement than:
%
$$
(A \cong B) \simeq (A \equiv B)
$$
%
Which is permissable simply by ``forgetting'' that $\idToIso_{A\ B}$ plays the
role of the equivalence.

With all this in place it is now possible to prove that all the laws are indeed
mere propositions. Most of the proofs simply use the fact that the type of
arrows are sets. This is because most of the laws are a collection of equations
between arrows in the category. And since such a proof does not have any
content, two witnesses must be the same. All the proofs are really quite
mechanical. Lets have a look at one of them: The identity law states that:
%
$$
\prod_{A\ B \tp \Object} \prod_{f \tp A \to B} \id \comp f \equiv f \x f \comp \id \equiv f
$$
%
There are multiple ways to prove this. Perhaps one of the more intuitive proofs
is by way of the following `combinators':
%
$$
\mathit{propPi} \tp \left(\prod_{a \tp A} \isProp\ (P\ a)\right) \to \isProp\ \left(\prod_{a \tp A} P\ a\right)
$$
%
I.e.; pi-types preserve propositionality when the co-domain is always a
proposition.
%
$$
\mathit{propSig} \tp \isProp\ A \to \left(\prod_{a \tp A} \isProp\ (P\ a)\right) \to \isProp\ \left(\sum_{a \tp A} P\ a\right)
$$
%
I.e.; sigma-types preserve propositionality whenever it's first component is a
proposition, and it's second component is always a proposition for all points of
in the left type.

So the proof goes like this: We `eliminate' the 3 function abstractions by
applying $\propPi$ three times, then we eliminate the (non-dependent) sigma-type
by applying $\propSig$ and are thus left with the two proof-obligations:
$\isProp\ (\id \comp f \equiv f)$ and $\isProp\ (f \comp \id \equiv f)$ which
follows from the type of arrows being a set.

This example illustrates nicely how we can use these combinators to reason about
`canonical' types like $\sum$ and $\prod$. Similiar combinators can be defined
at the other homotopic levels. These combinators are however not applicable in
situations where we want to reason about other types - e.g. types we've defined
ourselves. For instance, after we've proven that all the projections of
pre-categories are propositions, we would like to bundle this up to show that
the type of pre-categories is also a proposition. Since pre-categories are not
formulates with a chain of sigma-types we wont have any combinators available to
help us here. In stead we'll use the path-type directly.

What we want to prove is:
%
$$
\isProp\ \IsPreCategory
$$
%
Which is judgmentally the same as
%
$$
\prod_{a\ b \tp \IsPreCategory} a \equiv b
$$
%
So let $a\ b \tp \IsPreCategory$ be given. To prove the equality $a \equiv b$ is
to give a continuous path from the index-type into path-space - in this case
$\IsPreCategory$. This path must satisfy being being judgmentally the same as
$a$ at the left endpoint and $b$ at the right endpoint. I.e. a function $I \to
\IsPreCategory$. We know we can form a continuous path between all projections
of $a$ and $b$, this follows from the type of all the projections being mere
propositions. For instance, the path between $\isIdentity_a$ and $\isIdentity_b$
is simply formed by:
%
$$
\propIsIdentity\ \isIdentity_a\ \isIdentity_b \tp \isIdentity_a \equiv \isIdentity_b
$$
%
So to give the continuous function $I \to \IsPreCategory$ that is our goal we
introduce $i \tp I$ and proceed by constructing an element of $\IsPreCategory$
by using that all the projections are propositions to generate paths between all
projections. Once we have such a path e.g. $p : \isIdentity_a \equiv
\isIdentity_b$ we can elimiate it with $i$ and thus obtaining $p\ i \tp
\isIdentity_{p\ i}$ and this element satisfies exactly that it corresponds to
the corresponding projections at either endpoint. Thus the element we construct
at $i$ becomes:
%
\begin{align*}
  & \{\ \mathit{propIsAssociative}\   \mathit{isAssociative}_x\
       \mathit{isAssociative}_y\ i  \\
  & ,\  \mathit{propIsIdentity}\      \mathit{isIdentity}_x\
       \mathit{isIdentity}_y\ i  \\
  & ,\  \mathit{propArrowsAreSets}\ \mathit{arrowsAreSets}_x\
       \mathit{arrowsAreSets}_y\ i  \\
  & \}
\end{align*}
%
I've found that this to be a general pattern when proving things in homotopy
type theory, namely that you have to wrap and unwrap equalities at different
levels. It is worth noting that proving this theorem with the regular inductive
equality type would already not be possible, since we at least need
extensionality (the projections are all $\prod$-types). Assuming we had
functional extensionality available to us as an axiom, we would use functional
extensionality (in reverse?) to retreive the equalities in $a$ and $b$,
pattern-match on them to see that they are both $\mathit{refl}$ and then close
the proof with $\mathit{refl}$. Of course this theorem is not so interesting in
the setting of ITT since we know a priori that equality proofs are unique.

The situation is a bit more complicated when we have a dependent type. For
instance, when we want to show that $\IsCategory$ is a mere proposition.
$\IsCategory$ is a record with two fields, a witness to being a pre-category and
the univalence condition. Recall that the univalence condition is indexed by the
identity-proof. So if we follow the same recipe as above, let $a\ b \tp
\IsCategory$, to show them equal, we now need to give two paths. One homogenous:
%
$$
p_{\isPreCategory} \tp \isPreCategory_a \equiv \isPreCategory_b
$$
%
and one heterogeneous:
%
$$
Path\ (\Gl i \to Univalent_{p\ i})\ \isPreCategory_a\ \isPreCategory_b
$$
%
Which depends on the choice of $p_{\isPreCategory}$. The first of these we can
provide since, as we have shown, $\IsPreCategory$ is a proposition. However,
even though $\Univalent$ is also a proposition, we cannot use this directly to
show the latter. This is becasue $\isProp$ talks about non-dependent paths. To
`promote' this to a dependent path we can use another useful combinator;
$\lemPropF$. Given a type $A \tp \MCU$ and a type family on $A$; $B : A \to
\MCU$. Let $P$ be a proposition indexed by an element of $A$ and say we have a
path between some two elements in $A$; $p : a_0 \equiv a_1$ then we can built a
heterogeneous path between any two $b$'s at the endpoints:
%
$$
Path\ (\Gl i \to B\ (p\ i))\ b0\ b1
$$
%
where $b_0 \tp B a_0$ and $b_1 \tp B\ a_1$. This is quite a mouthful, but the
example at present should serve as an illustration. In this case $A =
\mathit{IsIdentity}\ \mathit{identity}$ and $B = \mathit{Univalent}$ we've shown
that being a category is a proposition, a result that holds for any choice of
identity proof. Finally we must provide a proof that the identity proofs at $a$
and $b$ are indeed the same, this we can extract from $p_{\isPreCategory}$ by
applying using congruence of paths: $\congruence\ \mathit{isIdentity}\
p_{\isPreCategory}$

When we have a proper category we can make precise the notion of ``identifying
isomorphic types'' (TODO cite awodey here). That is, we can construct the
function:
%
$$
\isoToId \tp (A \cong B) \to (A \equiv B)
$$
%
One application of this, and a perhaps somewhat surprising result, is that
terminal objects are propositional. Intuitively; they do not have any
interesting structure. The proof of this follows from the usual observation that
any two terminal objects are isomorphic. The proof is omitted here, but the
curious reader can check the implementation for the details. (TODO: The proof is
a bit fun, should I include it?)

In the following I will demonstrate how to instantiate a category and
subsequently why the result above is very useful to have when equating
categories (TODO: This promise is not fulfilled immediately as I digress and
talk about equivalences). So let us define the notion of the opposite category.
This is arguably one of the simplest constructions of a category one can give.
Let $\bC$ be a category, we then define a new category called the opposite of
$\bC$; $\overline{\bC}$. It has the same objects and the same identity, an arrow
from $A$ to $B$ in this category corresponds to an arrow from $B$ to $A$ in the
underlying category. Function composition will then be reverse function
composition from the underlying category.

Showing that this forms a pre-category is rather straightforward. I'll state the
laws in terms of the underlying category $\bC$:
%
$$
h >>> (g >>> f) \equiv h >>> g >>> f
$$
%
Since $>>>$ is reverse function composition this is just the symmetric version
of associativity.
%
$$
\matit{identity} >>> f \equiv f \x f >>> identity \equiv f
$$
This is just the swapped version of identity.

Finally, that the arrows form sets just follows by flipping the order of the
arguments. Or in other words since $\Hom_{A\ B}$ is a set for all $A\ B \tp
\Object$ then so is $\Hom_{B\ A}$.

Now, to show that this category is univalent is not as trivial. So I will
digress at this point and talk about equivalences. We will return to this category in section ????.

\subsection{Equivalences}
The usual notion of a function $f : A \to B$ having an inverses is:
%
$$
\sum_{g : B \to A} f \comp g \equiv \identity_{B} \x g \comp f \equiv \identity_{A}
$$
%
This is defined in \cite[p. 129]{hott-2013} and is referred to as the a
quasi-inverse. At the same place \cite{hott-2013} gives an ``interface'' for
what an equivalence $\isequiv : (A \to B) \to \MCU$ must supply:
%
\begin{itemize}
\item
  $\qinv\ f \to \isequiv\ f$
\item
  $\isequiv\ f \to \qinv\ f$
\item
  $\isequiv\ f$ is a proposition
\end{itemize}
%
Having such an interface us to both 1) think rather abstractly about how to work
with equivalences and 2) to use ad-hoc definitions of equivalences. The specific
instantiation of $\isequiv$ as defined in \cite{cubical} is:
%
$$
isEquiv\ f \defeq \prod_{b : B} \isContr\ (\fiber\ f\ b)
$$
where
$$
\fiber\ f\ b \defeq \sum_{a \tp A} b \equiv f\ a
$$
%
I give it's definition here mainly for completeness, because as I stated we can
move away from this specific instantiation and think about it more abstractly
once we have shown that this definition actually works as an equivalence.

The first function from the list of requirements we will call
$\mathit{fromIsomorphism}$, this is known as $\mathit{gradLemma}$ in
\cite{cubical} the second one we will refer to as $\mathit{toIsmorphism}$. It's
implementation can be found in the sources. Likewise the proof that this
equivalence is propositional can be found in my implementation.

So another way to provide a proof that a category is univalent is to give give
an inverse to $\idToIso\ A\ B$. I want to stress here that the notion of an
inverse at this point is conflated. There is the notion of an inverse in the
context of a category (where the concept of functions are generalized to arrows)
and, as here, an inverse as a regular type-theoretic function. This is
particularly confusing because the function that one must give the inverse to
has the type
%
$$
(A \cong B) \to (A \equiv B)
$$
%
where $\cong$ refers to ismorphism \emph{in the category}!

TODO: There is a lot more to say about equivalences!

\subsection{Categories contd.}
Back from this aside, we can now show that the opposite category is also
univalent simply by showing that $\idToIso \tp (A \equiv B) \to (A \cong B)$ is
an isomorphism (seen as a function). Dually we have that $\idToIso_{\bC} \tp (A
\equiv B) \to (A \cong_{\bC} B)$ is an isomorphism. Let us denote it's inverse
as $\eta \tp (A \cong_{\bC} B) \to (A \equiv B)$. If we squint we can see what
we need is a way to go between $\cong$ and $\cong_{\bC}$, well, an inhabitant of
$A \cong B$ is simply a pair of arrows $f$ being the isomorphism and $g$ it's
inverse. In the present category $g$ will play the role of the isomorphism and
$f$ will be the inverse. Similarly we can go in the opposite direction. These
two functions are obviously inverses. Name them $\mathit{shuffle} \tp (A \cong
B) \to (A \cong_{\bC} B)$ and $\mathit{shuffle}^{-1} : (A \cong_{\bC} B) \to (A
\cong B)$ respectively.

As the inverse of $\idToIso$ we will pick $\zeta \defeq \eta \comp
\mathit{shuffle}$. The proof that they are inverses go as follows:
%
\begin{align*}
\zeta \comp \idToIso & \equiv
\eta \comp \shuffle \comp \idToIso
&& \text{by definition} \\
%% ≡⟨ cong (λ φ → φ x) (cong (λ φ → η ⊙ shuffle ⊙ φ) (funExt lem)) ⟩ \\
%
& \equiv
\eta \comp \shuffle \comp \inv{\shuffle} \comp \idToIso
&& \text{lemma} \\
%%   ≡⟨⟩ \\
& \equiv
\eta \comp \idToIso_{\bC}
&& \text{$\shuffle$ is an isomorphism} \\
%%   ≡⟨ (λ i → verso-recto i x) ⟩ \\
& \equiv
\identity
&& \text{$\eta$ is an ismorphism} \\
\end{align*}
%
The other direction is analogous.

The lemma used in this proof show that $\idToIso \equiv \inv{\shuffle} \comp
\idToIso_{\bC}$ it's a rather straight-forward proof since being-an-inverse-of
is a proposition.

So, in conclusion, we've shown that the opposite category is indeed a category.
We can now proceed to show that this construction is an involution, namely:
%
$$
\prod_{\bC : \Category} \left(\bC^T\right)^T \equiv \bC
$$
%
As we've seen the laws in $\left(\bC^T\right)^T$ get quite involved.\footnote{We
  haven't even seen the full story because we've used this `interface' for
  equivalences.} Luckily they being a category is a proposition, so we need not
concern ourselves with this bit when proving the above. We can use the equality
principle for categories that lets us prove an equality just by giving an
equality on the data-part. So, given a category $\bC$ what we must provide is
the following proof:
%
$$
\mathit{raw}\ \left(\bC^T\right)^T \equiv \mathit{raw}\ \bC
$$
%
And these are judgmentally the same. I remind the reader that the left-hand side
is constructed by flipping the arrows, an action that is certainly an
involution.

%% \subsubsection{Functors}
%% Defines the notion of a functor - also split up into data and laws.

%% Propositionality for being a functor.

%% Composition of functors and the identity functor.

%% \subsubsection{Products}
%% Definition of what it means for an object to be a product in a given category.

%% Definition of what it means for a category to have all products.

%% \WIP{} Prove propositionality for being a product and having products.

%% \subsubsection{Exponentials}
%% Definition of what it means to be an exponential object.

%% Definition of what it means for a category to have all exponential objects.

%% \subsubsection{Cartesian closed categories}
%% Definition of what it means for a category to be cartesian closed; namely that
%% it has all products and all exponentials.

%% \subsubsection{Natural transformations}
%% Definition of transformations\footnote{Maybe this is a name I made up for a
%%   family of morphisms} and the naturality condition for these.

%% Proof that naturality is a mere proposition and the accompanying equality
%% principle. Proof that natural transformations are homotopic sets.

%% The identity natural transformation.

%% \subsubsection{Yoneda embedding}

%% The yoneda embedding is typically presented in terms of the category of
%% categories (cf. Awodey) \emph however this is not stricly needed - all we need
%% is what would be the exponential object in that category - this happens to be
%% functors and so this is how we define the yoneda embedding.

%% \subsubsection{Monads}

%% Defines an equivalence between these two formulations of a monad:

%% \subsubsubsection{Monoidal monads}

%% Defines the standard monoidal representation of a monad:

%% An endofunctor with two natural transformations (called ``pure'' and ``join'')
%% and some laws about these natural transformations.

%% Propositionality proofs and equality principle is provided.

%% \subsubsubsection{Kleisli monads}

%% A presentation of monads perhaps more familiar to a functional programer:

%% A map on objects and two maps on morphisms (called ``pure'' and ``bind'') and
%% some laws about these maps.

%% Propositionality proofs and equality principle is provided.

%% \subsubsubsection{Voevodsky's construction}

%% Provides construction 2.3 as presented in an unpublished paper by Vladimir
%% Voevodsky. This construction is similiar to the equivalence provided for the two
%% preceding formulations
%% \footnote{ TODO: I would like to include in the thesis some motivation for why
%%   this construction is particularly interesting.}

%% \subsubsection{Homotopy sets}
%% The typical category of sets where the objects are modelled by an Agda set
%% (henceforth ``$\Type$'') at a given level is not a valid category in this cubical
%% settings, we need to restrict the types to be those that are homotopy sets. Thus
%% the objects of this category are:
%% %
%% $$\hSet_\ell \defeq \sum_{A \tp \MCU_\ell} \isSet\ A$$
%% %
%% The definition of univalence for categories I have defined is:
%% %
%% $$\isEquiv\ (\hA \equiv \hB)\ (\hA \cong \hB)\ \idToIso$$
%% %
%% Where $\hA and \hB$ denote objects in the category. Note that this is stronger
%% than
%% %
%% $$(\hA \equiv \hB) \simeq (\hA \cong \hB)$$
%% %
%% Because we require that the equivalence is constructed from the witness to:
%% %
%% $$\id \comp f \equiv f \x f \comp \id \equiv f$$
%% %
%% And indeed univalence does not follow immediately from univalence for types:
%% %
%% $$(A \equiv B) \simeq (A \simeq B)$$
%% %
%% Because $A\ B \tp \Type$ whereas $\hA\ \hB \tp \hSet$.

%% For this reason I have shown that this category satisfies the following
%% equivalent formulation of being univalent:
%% %
%% $$\prod_{A \tp hSet} \isContr \left( \sum_{X \tp hSet} A \cong X \right)$$
%% %
%% But I have not shown that it is indeed equivalent to my former definition.
%% \subsubsection{Categories}
%% Note that this category does in fact not exist. In stead I provide the
%% definition of the ``raw'' category as well as some of the laws.

%% Furthermore I provide some helpful lemmas about this raw category. For instance
%% I have shown what would be the exponential object in such a category.

%% These lemmas can be used to provide the actual exponential object in a context
%% where we have a witness to this being a category. This is useful if this library
%% is later extended to talk about higher categories.

%% \subsubsection{Functors}
%% The category of functors and natural transformations. An immediate corrolary is
%% the set of presheaf categories.

%% \WIP{} I have not shown that the category of functors is univalent.

%% \subsubsection{Relations}
%% The category of relations. \WIP{} I have not shown that this category is
%% univalent. Not sure I intend to do so either.

%% \subsubsection{Free category}
%% The free category of a category. \WIP{} I have not shown that this category is
%% univalent.

%% \subsection{Current Challenges}
%% Besides the items marked \WIP{} above I still feel a bit unsure about what to
%% include in my report. Most of my work so far has been specifically about
%% developing this library. Some ideas:
%% %
%% \begin{itemize}
%% \item
%%   Modularity properties
%% \item
%%   Compare with setoid-approach to solve similiar problems.
%% \item
%%   How to structure an implementation to best deal with types that have no
%%   structure (propositions) and those that do (sets and everything above)
%% \end{itemize}
%% %
%% \subsection{Ideas for future developments}
%% \subsubsection{Higher categories}
%% I only have a notion of (1-)categories. Perhaps it would be nice to also
%% formalize higher categories.

%% \subsubsection{Hierarchy of concepts related to monads}
%% In Haskell the type-class Monad sits in a hierarchy atop the notion of a functor
%% and applicative functors. There's probably a similiar notion in the
%% category-theoretic approach to developing this.

%% As I have already defined monads from these two perspectives, it would be
%% interesting to take this idea even further and actually show how monads are
%% related to applicative functors and functors. I'm not entirely sure how this
%% would look in Agda though.

%% \subsubsection{Use formulation on the standard library}
%% I also thought it would be interesting to use this library to show certain
%% properties about functors, applicative functors and monads used in the Agda
%% Standard library. So I went ahead and tried to show that agda's standard
%% library's notion of a functor (along with suitable laws) is equivalent to my
%% formulation (in the category of homotopic sets). I ran into two problems here,
%% however; the first one is that the standard library's notion of a functor is
%% indexed by the object map:
%% %
%% $$
%% \Functor \tp (\Type \to \Type) \to \Type
%% $$
%% %
%% Where $\Functor\ F$ has the member:
%% %
%% $$
%% \fmap \tp (A \to B) \to F A \to F B
%% $$
%% %
%% Whereas the object map in my definition is existentially quantified:
%% %
%% $$
%% \Functor \tp \Type
%% $$
%% %
%% And $\Functor$ has these members:
%% \begin{align*}
%% F     & \tp \Type \to \Type \\
%% \fmap & \tp (A \to B) \to F A \to F B\}
%% \end{align*}
%% %